Information Theory, Inference, and Learning - Inference Group

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Copyright Cambridge University Press 2003. On-screen viewing permitted. Printing not permitted. http://www.cambridge.org/0521642981 You can buy this book for 30 pounds or $50. See http://www.inference.phy.cam.ac.uk/mackay/itila/ for links.

Information Theory, Inference, and Learning Algorithms

David J.C. MacKay

Copyright Cambridge University Press 2003. On-screen viewing permitted. Printing not permitted. http://www.cambridge.org/0521642981 You can buy this book for 30 pounds or $50. See http://www.inference.phy.cam.ac.uk/mackay/itila/ for links.

Information Theory, Inference, and Learning Algorithms

David J.C. MacKay [email protected] c

1995, 1996, 1997, 1998, 1999, 2000, 2001, 2002, 2003, 2004, 2005 c

Cambridge University Press 2003

Version 7.2 (fourth printing) March 28, 2005 Please send feedback on this book via http://www.inference.phy.cam.ac.uk/mackay/itila/ Version 6.0 of this book was published by C.U.P. in September 2003. It will remain viewable on-screen on the above website, in postscript, djvu, and pdf formats. In the second printing (version 6.6) minor typos were corrected, and the book design was slightly altered to modify the placement of section numbers. In the third printing (version 7.0) minor typos were corrected, and chapter 8 was renamed ‘Dependent random variables’ (instead of ‘Correlated’). In the fourth printing (version 7.2) minor typos were corrected. (C.U.P. replace this page with their own page ii.)

Copyright Cambridge University Press 2003. On-screen viewing permitted. Printing not permitted. http://www.cambridge.org/0521642981 You can buy this book for 30 pounds or $50. See http://www.inference.phy.cam.ac.uk/mackay/itila/ for links.

Contents

1 2 3

I

Data Compression 4 5 6 7

II

IV

. . . .

v 3 22 48

. . . . . . . . . . . . . . . . . . . . . .

65

The Source Coding Theorem . . Symbol Codes . . . . . . . . . . Stream Codes . . . . . . . . . . . Codes for Integers . . . . . . . .

Noisy-Channel Coding 8 9 10 11

III

Preface . . . . . . . . . . . . . . . . . . . Introduction to Information Theory . . . Probability, Entropy, and Inference . . . . More about Inference . . . . . . . . . . .

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137

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67 91 110 132

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Dependent Random Variables . . . . . . . . . . . Communication over a Noisy Channel . . . . . . The Noisy-Channel Coding Theorem . . . . . . . Error-Correcting Codes and Real Channels . . .

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138 146 162 177

Further Topics in Information Theory . . . . . . . . . . . . .

191

12 13 14 15 16 17 18 19

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193 206 229 233 241 248 260 269

Probabilities and Inference . . . . . . . . . . . . . . . . . .

281

20 21 22 23 24 25 26 27

284 293 300 311 319 324 334 341

Hash Codes: Codes for Efficient Information Retrieval Binary Codes . . . . . . . . . . . . . . . . . . . . . . Very Good Linear Codes Exist . . . . . . . . . . . . . Further Exercises on Information Theory . . . . . . . Message Passing . . . . . . . . . . . . . . . . . . . . . Communication over Constrained Noiseless Channels Crosswords and Codebreaking . . . . . . . . . . . . . Why have Sex? Information Acquisition and Evolution

An Example Inference Task: Clustering . Exact Inference by Complete Enumeration Maximum Likelihood and Clustering . . . Useful Probability Distributions . . . . . Exact Marginalization . . . . . . . . . . . Exact Marginalization in Trellises . . . . Exact Marginalization in Graphs . . . . . Laplace’s Method . . . . . . . . . . . . .

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Copyright Cambridge University Press 2003. On-screen viewing permitted. Printing not permitted. http://www.cambridge.org/0521642981 You can buy this book for 30 pounds or $50. See http://www.inference.phy.cam.ac.uk/mackay/itila/ for links.

28 29 30 31 32 33 34

Model Comparison and Occam’s Razor . . . . . . . . . . . Monte Carlo Methods . . . . . . . . . . . . . . . . . . . . . Efficient Monte Carlo Methods . . . . . . . . . . . . . . . . Ising Models . . . . . . . . . . . . . . . . . . . . . . . . . . Exact Monte Carlo Sampling . . . . . . . . . . . . . . . . . Variational Methods . . . . . . . . . . . . . . . . . . . . . . Independent Component Analysis and Latent Variable Modelling . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . Random Inference Topics . . . . . . . . . . . . . . . . . . . Decision Theory . . . . . . . . . . . . . . . . . . . . . . . . Bayesian Inference and Sampling Theory . . . . . . . . . .

437 445 451 457

Neural networks . . . . . . . . . . . . . . . . . . . . . . . .

467

38 39 40 41 42 43 44 45 46

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468 471 483 492 505 522 527 535 549

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555

35 36 37 V

VI

Introduction to Neural Networks The Single Neuron as a Classifier Capacity of a Single Neuron . . . Learning as Inference . . . . . . Hopfield Networks . . . . . . . . Boltzmann Machines . . . . . . . Supervised Learning in Multilayer Gaussian Processes . . . . . . . Deconvolution . . . . . . . . . .

Sparse Graph Codes 47 48 49 50

. . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . Networks . . . . . . . . . . . .

Low-Density Parity-Check Codes . . Convolutional Codes and Turbo Codes Repeat–Accumulate Codes . . . . . . Digital Fountain Codes . . . . . . . .

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597

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VII Appendices . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . .

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557 574 582 589

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. . . .

A Notation . . . . . B Some Physics . . . C Some Mathematics Bibliography . . . . . . Index . . . . . . . . . .

. . . .

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343 357 387 400 413 422

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598 601 605 613 620

Copyright Cambridge University Press 2003. On-screen viewing permitted. Printing not permitted. http://www.cambridge.org/0521642981 You can buy this book for 30 pounds or $50. See http://www.inference.phy.cam.ac.uk/mackay/itila/ for links.

Preface This book is aimed at senior undergraduates and graduate students in Engineering, Science, Mathematics, and Computing. It expects familiarity with calculus, probability theory, and linear algebra as taught in a first- or secondyear undergraduate course on mathematics for scientists and engineers. Conventional courses on information theory cover not only the beautiful theoretical ideas of Shannon, but also practical solutions to communication problems. This book goes further, bringing in Bayesian data modelling, Monte Carlo methods, variational methods, clustering algorithms, and neural networks. Why unify information theory and machine learning? Because they are two sides of the same coin. In the 1960s, a single field, cybernetics, was populated by information theorists, computer scientists, and neuroscientists, all studying common problems. Information theory and machine learning still belong together. Brains are the ultimate compression and communication systems. And the state-of-the-art algorithms for both data compression and error-correcting codes use the same tools as machine learning.

How to use this book The essential dependencies between chapters are indicated in the figure on the next page. An arrow from one chapter to another indicates that the second chapter requires some of the first. Within Parts I, II, IV, and V of this book, chapters on advanced or optional topics are towards the end. All chapters of Part III are optional on a first reading, except perhaps for Chapter 16 (Message Passing). The same system sometimes applies within a chapter: the final sections often deal with advanced topics that can be skipped on a first reading. For example in two key chapters – Chapter 4 (The Source Coding Theorem) and Chapter 10 (The Noisy-Channel Coding Theorem) – the first-time reader should detour at section 4.5 and section 10.4 respectively. Pages vii–x show a few ways to use this book. First, I give the roadmap for a course that I teach in Cambridge: ‘Information theory, pattern recognition, and neural networks’. The book is also intended as a textbook for traditional courses in information theory. The second roadmap shows the chapters for an introductory information theory course and the third for a course aimed at an understanding of state-of-the-art error-correcting codes. The fourth roadmap shows how to use the text in a conventional course on machine learning.

v

Copyright Cambridge University Press 2003. On-screen viewing permitted. Printing not permitted. http://www.cambridge.org/0521642981 You can buy this book for 30 pounds or $50. See http://www.inference.phy.cam.ac.uk/mackay/itila/ for links.

vi

Preface 1

Introduction to Information Theory

IV Probabilities and Inference

2

Probability, Entropy, and Inference

20

An Example Inference Task: Clustering

3

More about Inference

21

Exact Inference by Complete Enumeration

22

Maximum Likelihood and Clustering

23

Useful Probability Distributions

I

Data Compression

4

The Source Coding Theorem

24

Exact Marginalization

5

Symbol Codes

25

Exact Marginalization in Trellises

6

Stream Codes

26

Exact Marginalization in Graphs

7

Codes for Integers

27

Laplace’s Method

28

Model Comparison and Occam’s Razor

29

Monte Carlo Methods

II

Noisy-Channel Coding

8

Dependent Random Variables

30

Efficient Monte Carlo Methods

9

Communication over a Noisy Channel

31

Ising Models

10

The Noisy-Channel Coding Theorem

32

Exact Monte Carlo Sampling

11

Error-Correcting Codes and Real Channels

33

Variational Methods

34

Independent Component Analysis

35

Random Inference Topics

III Further Topics in Information Theory 12

Hash Codes

36

Decision Theory

13

Binary Codes

37

Bayesian Inference and Sampling Theory

14

Very Good Linear Codes Exist

15

Further Exercises on Information Theory

16

Message Passing

38

Introduction to Neural Networks

17

Constrained Noiseless Channels

39

The Single Neuron as a Classifier

18

Crosswords and Codebreaking

40

Capacity of a Single Neuron

19

Why have Sex?

41

Learning as Inference

42

Hopfield Networks

43

Boltzmann Machines

44

Supervised Learning in Multilayer Networks

45

Gaussian Processes

46

Deconvolution

V

Neural networks

VI Sparse Graph Codes

Dependencies

47

Low-Density Parity-Check Codes

48

Convolutional Codes and Turbo Codes

49

Repeat–Accumulate Codes

50

Digital Fountain Codes

Copyright Cambridge University Press 2003. On-screen viewing permitted. Printing not permitted. http://www.cambridge.org/0521642981 You can buy this book for 30 pounds or $50. See http://www.inference.phy.cam.ac.uk/mackay/itila/ for links.

Preface

vii

1

Introduction to Information Theory

IV Probabilities and Inference

2

Probability, Entropy, and Inference

20

An Example Inference Task: Clustering

3

More about Inference

21

Exact Inference by Complete Enumeration

22

Maximum Likelihood and Clustering

23

Useful Probability Distributions

I

Data Compression

4

The Source Coding Theorem

24

Exact Marginalization

5

Symbol Codes

25

Exact Marginalization in Trellises

6

Stream Codes

26

Exact Marginalization in Graphs

7

Codes for Integers

27

Laplace’s Method

28

Model Comparison and Occam’s Razor

29

Monte Carlo Methods

II

Noisy-Channel Coding

8

Dependent Random Variables

30

Efficient Monte Carlo Methods

9

Communication over a Noisy Channel

31

Ising Models

10

The Noisy-Channel Coding Theorem

32

Exact Monte Carlo Sampling

11

Error-Correcting Codes and Real Channels

33

Variational Methods

34

Independent Component Analysis

35

Random Inference Topics

III Further Topics in Information Theory 12

Hash Codes

36

Decision Theory

13

Binary Codes

37

Bayesian Inference and Sampling Theory

14

Very Good Linear Codes Exist

15

Further Exercises on Information Theory

16

Message Passing

38

Introduction to Neural Networks

17

Constrained Noiseless Channels

39

The Single Neuron as a Classifier

18

Crosswords and Codebreaking

40

Capacity of a Single Neuron

19

Why have Sex?

41

Learning as Inference

42

Hopfield Networks

43

Boltzmann Machines

44

Supervised Learning in Multilayer Networks

45

Gaussian Processes

46

Deconvolution

My Cambridge Course on, Information Theory, Pattern Recognition, and Neural Networks

V

Neural networks

VI Sparse Graph Codes 47

Low-Density Parity-Check Codes

48

Convolutional Codes and Turbo Codes

49

Repeat–Accumulate Codes

50

Digital Fountain Codes

Copyright Cambridge University Press 2003. On-screen viewing permitted. Printing not permitted. http://www.cambridge.org/0521642981 You can buy this book for 30 pounds or $50. See http://www.inference.phy.cam.ac.uk/mackay/itila/ for links.

viii

Preface 1

Introduction to Information Theory

IV Probabilities and Inference

2

Probability, Entropy, and Inference

20

An Example Inference Task: Clustering

3

More about Inference

21

Exact Inference by Complete Enumeration

22

Maximum Likelihood and Clustering

23

Useful Probability Distributions

I

Data Compression

4

The Source Coding Theorem

24

Exact Marginalization

5

Symbol Codes

25

Exact Marginalization in Trellises

6

Stream Codes

26

Exact Marginalization in Graphs

7

Codes for Integers

27

Laplace’s Method

28

Model Comparison and Occam’s Razor

29

Monte Carlo Methods

II

Noisy-Channel Coding

8

Dependent Random Variables

30

Efficient Monte Carlo Methods

9

Communication over a Noisy Channel

31

Ising Models

10

The Noisy-Channel Coding Theorem

32

Exact Monte Carlo Sampling

11

Error-Correcting Codes and Real Channels

33

Variational Methods

34

Independent Component Analysis

35

Random Inference Topics

III Further Topics in Information Theory 12

Hash Codes

36

Decision Theory

13

Binary Codes

37

Bayesian Inference and Sampling Theory

14

Very Good Linear Codes Exist

15

Further Exercises on Information Theory

16

Message Passing

38

Introduction to Neural Networks

17

Constrained Noiseless Channels

39

The Single Neuron as a Classifier

18

Crosswords and Codebreaking

40

Capacity of a Single Neuron

19

Why have Sex?

41

Learning as Inference

42

Hopfield Networks

43

Boltzmann Machines

44

Supervised Learning in Multilayer Networks

45

Gaussian Processes

46

Deconvolution

V

Neural networks

VI Sparse Graph Codes Short Course on Information Theory

47

Low-Density Parity-Check Codes

48

Convolutional Codes and Turbo Codes

49

Repeat–Accumulate Codes

50

Digital Fountain Codes

Copyright Cambridge University Press 2003. On-screen viewing permitted. Printing not permitted. http://www.cambridge.org/0521642981 You can buy this book for 30 pounds or $50. See http://www.inference.phy.cam.ac.uk/mackay/itila/ for links.

Preface

ix

1

Introduction to Information Theory

IV Probabilities and Inference

2

Probability, Entropy, and Inference

20

An Example Inference Task: Clustering

3

More about Inference

21

Exact Inference by Complete Enumeration

22

Maximum Likelihood and Clustering

23

Useful Probability Distributions

I

Data Compression

4

The Source Coding Theorem

24

Exact Marginalization

5

Symbol Codes

25

Exact Marginalization in Trellises

6

Stream Codes

26

Exact Marginalization in Graphs

7

Codes for Integers

27

Laplace’s Method

28

Model Comparison and Occam’s Razor

29

Monte Carlo Methods

II

Noisy-Channel Coding

8

Dependent Random Variables

30

Efficient Monte Carlo Methods

9

Communication over a Noisy Channel

31

Ising Models

10

The Noisy-Channel Coding Theorem

32

Exact Monte Carlo Sampling

11

Error-Correcting Codes and Real Channels

33

Variational Methods

34

Independent Component Analysis

35

Random Inference Topics

III Further Topics in Information Theory 12

Hash Codes

36

Decision Theory

13

Binary Codes

37

Bayesian Inference and Sampling Theory

14

Very Good Linear Codes Exist

15

Further Exercises on Information Theory

16

Message Passing

38

Introduction to Neural Networks

17

Constrained Noiseless Channels

39

The Single Neuron as a Classifier

18

Crosswords and Codebreaking

40

Capacity of a Single Neuron

19

Why have Sex?

41

Learning as Inference

42

Hopfield Networks

43

Boltzmann Machines

44

Supervised Learning in Multilayer Networks

45

Gaussian Processes

46

Deconvolution

V

Neural networks

VI Sparse Graph Codes Advanced Course on Information Theory and Coding

47

Low-Density Parity-Check Codes

48

Convolutional Codes and Turbo Codes

49

Repeat–Accumulate Codes

50

Digital Fountain Codes

Copyright Cambridge University Press 2003. On-screen viewing permitted. Printing not permitted. http://www.cambridge.org/0521642981 You can buy this book for 30 pounds or $50. See http://www.inference.phy.cam.ac.uk/mackay/itila/ for links.

x

Preface 1

Introduction to Information Theory

IV Probabilities and Inference

2

Probability, Entropy, and Inference

20

An Example Inference Task: Clustering

3

More about Inference

21

Exact Inference by Complete Enumeration

22

Maximum Likelihood and Clustering

23

Useful Probability Distributions

I

Data Compression

4

The Source Coding Theorem

24

Exact Marginalization

5

Symbol Codes

25

Exact Marginalization in Trellises

6

Stream Codes

26

Exact Marginalization in Graphs

7

Codes for Integers

27

Laplace’s Method

28

Model Comparison and Occam’s Razor

29

Monte Carlo Methods

II

Noisy-Channel Coding

8

Dependent Random Variables

30

Efficient Monte Carlo Methods

9

Communication over a Noisy Channel

31

Ising Models

10

The Noisy-Channel Coding Theorem

32

Exact Monte Carlo Sampling

11

Error-Correcting Codes and Real Channels

33

Variational Methods

34

Independent Component Analysis

35

Random Inference Topics

III Further Topics in Information Theory 12

Hash Codes

36

Decision Theory

13

Binary Codes

37

Bayesian Inference and Sampling Theory

14

Very Good Linear Codes Exist

15

Further Exercises on Information Theory

16

Message Passing

38

Introduction to Neural Networks

17

Constrained Noiseless Channels

39

The Single Neuron as a Classifier

18

Crosswords and Codebreaking

40

Capacity of a Single Neuron

19

Why have Sex?

41

Learning as Inference

42

Hopfield Networks

43

Boltzmann Machines

44

Supervised Learning in Multilayer Networks

45

Gaussian Processes

46

Deconvolution

V

Neural networks

VI Sparse Graph Codes A Course on Bayesian Inference and Machine Learning

47

Low-Density Parity-Check Codes

48

Convolutional Codes and Turbo Codes

49

Repeat–Accumulate Codes

50

Digital Fountain Codes

Copyright Cambridge University Press 2003. On-screen viewing permitted. Printing not permitted. http://www.cambridge.org/0521642981 You can buy this book for 30 pounds or $50. See http://www.inference.phy.cam.ac.uk/mackay/itila/ for links.

Preface

xi

About the exercises You can understand a subject only by creating it for yourself. The exercises play an essential role in this book. For guidance, each has a rating (similar to that used by Knuth (1968)) from 1 to 5 to indicate its difficulty. In addition, exercises that are especially recommended are marked by a marginal encouraging rat. Some exercises that require the use of a computer are marked with a C. Answers to many exercises are provided. Use them wisely. Where a solution is provided, this is indicated by including its page number alongside the difficulty rating. Solutions to many of the other exercises will be supplied to instructors using this book in their teaching; please email [email protected]

Summary of codes for exercises Especially recommended . C [p. 42]

Recommended Parts require a computer Solution provided on page 42

[1 ] [2 ] [3 ] [4 ] [5 ]

Simple (one minute) Medium (quarter hour) Moderately hard Hard Research project

Internet resources The website http://www.inference.phy.cam.ac.uk/mackay/itila contains several resources: 1. Software. Teaching software that I use in lectures, interactive software, and research software, written in perl, octave, tcl, C, and gnuplot. Also some animations. 2. Corrections to the book. Thank you in advance for emailing these! 3. This book. The book is provided in postscript, pdf, and djvu formats for on-screen viewing. The same copyright restrictions apply as to a normal book.

About this edition This is the fourth printing of the first edition. In the second printing, the design of the book was altered slightly. Page-numbering generally remained unchanged, except in chapters 1, 6, and 28, where a few paragraphs, figures, and equations moved around. All equation, section, and exercise numbers were unchanged. In the third printing, chapter 8 was renamed ‘Dependent Random Variables’, instead of ‘Correlated’, which was sloppy.

Copyright Cambridge University Press 2003. On-screen viewing permitted. Printing not permitted. http://www.cambridge.org/0521642981 You can buy this book for 30 pounds or $50. See http://www.inference.phy.cam.ac.uk/mackay/itila/ for links.

xii

Preface

Acknowledgments I am most grateful to the organizations who have supported me while this book gestated: the Royal Society and Darwin College who gave me a fantastic research fellowship in the early years; the University of Cambridge; the Keck Centre at the University of California in San Francisco, where I spent a productive sabbatical; and the Gatsby Charitable Foundation, whose support gave me the freedom to break out of the Escher staircase that book-writing had become. My work has depended on the generosity of free software authors. I wrote the book in LATEX 2ε . Three cheers for Donald Knuth and Leslie Lamport! Our computers run the GNU/Linux operating system. I use emacs, perl, and gnuplot every day. Thank you Richard Stallman, thank you Linus Torvalds, thank you everyone. Many readers, too numerous to name here, have given feedback on the book, and to them all I extend my sincere acknowledgments. I especially wish to thank all the students and colleagues at Cambridge University who have attended my lectures on information theory and machine learning over the last nine years. The members of the Inference research group have given immense support, and I thank them all for their generosity and patience over the last ten years: Mark Gibbs, Michelle Povinelli, Simon Wilson, Coryn Bailer-Jones, Matthew Davey, Katriona Macphee, James Miskin, David Ward, Edward Ratzer, Seb Wills, John Barry, John Winn, Phil Cowans, Hanna Wallach, Matthew Garrett, and especially Sanjoy Mahajan. Thank you too to Graeme Mitchison, Mike Cates, and Davin Yap. Finally I would like to express my debt to my personal heroes, the mentors from whom I have learned so much: Yaser Abu-Mostafa, Andrew Blake, John Bridle, Peter Cheeseman, Steve Gull, Geoff Hinton, John Hopfield, Steve Luttrell, Robert MacKay, Bob McEliece, Radford Neal, Roger Sewell, and John Skilling.

Dedication This book is dedicated to the campaign against the arms trade. www.caat.org.uk

Peace cannot be kept by force. It can only be achieved through understanding. – Albert Einstein

Copyright Cambridge University Press 2003. On-screen viewing permitted. Printing not permitted. http://www.cambridge.org/0521642981 You can buy this book for 30 pounds or $50. See http://www.inference.phy.cam.ac.uk/mackay/itila/ for links.

About Chapter 1 In the first chapter, you will need to be familiar with the binomial distribution. And to solve the exercises in the text – which I urge you to do – you will need to know Stirling’s approximation for the factorial function, x! ' x x e−x , and  N! N be able to apply it to r = (N −r)! r! . These topics are reviewed below.

Unfamiliar notation? See Appendix A, p.598.

The binomial distribution Example 1.1. A bent coin has probability f of coming up heads. The coin is tossed N times. What is the probability distribution of the number of heads, r? What are the mean and variance of r? Solution. The number of heads has a binomial distribution.   N r f (1 − f )N −r . P (r | f, N ) = r

(1.1)

0 1 2 3 4 5 6 7 8 9 10

The mean, E[r], and variance, var[r], of this distribution are defined by E[r] ≡

N X r=0

P (r | f, N ) r

i h var[r] ≡ E (r − E[r])2

r

(1.2)

(1.3)

= E[r 2 ] − (E[r])2 =

N X r=0

P (r | f, N )r 2 − (E[r])2 .

(1.4)

Rather than evaluating the sums over r in (1.2) and (1.4) directly, it is easiest to obtain the mean and variance by noting that r is the sum of N independent random variables, namely, the number of heads in the first toss (which is either zero or one), the number of heads in the second toss, and so forth. In general, E[x + y] = E[x] + E[y] for any random variables x and y; var[x + y] = var[x] + var[y] if x and y are independent.

(1.5)

So the mean of r is the sum of the means of those random variables, and the variance of r is the sum of their variances. The mean number of heads in a single toss is f × 1 + (1 − f ) × 0 = f , and the variance of the number of heads in a single toss is 

 f × 12 + (1 − f ) × 02 − f 2 = f − f 2 = f (1 − f ),

(1.6)

so the mean and variance of r are: E[r] = N f

var[r] = N f (1 − f ).

and 1

2

0.3 0.25 0.2 0.15 0.1 0.05 0

(1.7)

Figure 1.1. The binomial distribution P (r | f = 0.3, N = 10).

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2

About Chapter 1

Approximating x! and

N r



Let’s derive Stirling’s approximation by an unconventional route. We start from the Poisson distribution with mean λ,

0.12 0.1 0.08 0.06

P (r | λ) = e

−λ λ

r

r!

0.04

r ∈ {0, 1, 2, . . .}.

(1.8)

0.02 0 0

For large λ, this distribution is well approximated – at least in the vicinity of r ' λ – by a Gaussian distribution with mean λ and variance λ: e−λ

λr

1 ' √ r! 2πλ

(r−λ)2 − 2λ e

.

(1.9)

5

10

15

20

25

r Figure 1.2. The Poisson distribution P (r | λ = 15).

Let’s plug r = λ into this formula, then rearrange it. λλ λ!

1 √ 2πλ √ ⇒ λ! ' λλ e−λ 2πλ.

e−λ

'

This is Stirling’s approximation for the factorial function. √ x! ' xx e−x 2πx ⇔ ln x! ' x ln x − x + 12 ln 2πx.

(1.10) (1.11)

(1.12)

We have derived not only the leading order behaviour, x! ' x x e−x , but also, √ at no cost, the next-order correction term 2πx. We now apply Stirling’s  approximation to ln Nr :   N N N N! ' (N − r) ln + r ln . (1.13) ln ≡ ln (N − r)! r! N −r r r Since all the terms in this equation are logarithms, this result can be rewritten in any base. We will denote natural logarithms (log e ) by ‘ln’, and logarithms to base 2 (log 2 ) by ‘log’. If we introduce the binary entropy function, H2 (x) ≡ x log

1 1 + (1−x) log , x (1−x)

then we can rewrite the approximation (1.13) as   N log ' N H2 (r/N ), r or, equivalently,

  N ' 2N H2 (r/N ) . r

loge x . loge 2 ∂ log2 x 1 1 Note that = . ∂x loge 2 x

Recall that log2 x =

(1.14)

H2 (x)

1

0.8

(1.15)

0.6 0.4 0.2

(1.16)

If we need a more accurate approximation, we can include terms of the next order from Stirling’s approximation (1.12):     N N −r r ' N H2 (r/N ) − 12 log 2πN . (1.17) log r N N

0 0

0.2

0.4

0.6

0.8

1

x

Figure 1.3. The binary entropy function.

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1 Introduction to Information Theory The fundamental problem of communication is that of reproducing at one point either exactly or approximately a message selected at another point. (Claude Shannon, 1948) In the first half of this book we study how to measure information content; we learn how to compress data; and we learn how to communicate perfectly over imperfect communication channels. We start by getting a feeling for this last problem.

1.1 How can we achieve perfect communication over an imperfect, noisy communication channel? Some examples of noisy communication channels are: modem

- phone -

Galileo

- radio -

• an analogue telephone line, over which two modems communicate digital information; • the radio communication link from Galileo, the Jupiter-orbiting spacecraft, to earth; • reproducing cells, in which the daughter cells’ DNA contains information from the parent cells; • a disk drive. The last example shows that communication doesn’t have to involve information going from one place to another. When we write a file on a disk drive, we’ll read it off in the same location – but at a later time. These channels are noisy. A telephone line suffers from cross-talk with other lines; the hardware in the line distorts and adds noise to the transmitted signal. The deep space network that listens to Galileo’s puny transmitter receives background radiation from terrestrial and cosmic sources. DNA is subject to mutations and damage. A disk drive, which writes a binary digit (a one or zero, also known as a bit) by aligning a patch of magnetic material in one of two orientations, may later fail to read out the stored binary digit: the patch of material might spontaneously flip magnetization, or a glitch of background noise might cause the reading circuit to report the wrong value for the binary digit, or the writing head might not induce the magnetization in the first place because of interference from neighbouring bits. In all these cases, if we transmit data, e.g., a string of bits, over the channel, there is some probability that the received message will not be identical to the 3

line

waves

modem

Earth

daughter cell parent  cell @ R @ daughter cell computer - disk - computer memory memory drive

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4

1 — Introduction to Information Theory

transmitted message. We would prefer to have a communication channel for which this probability was zero – or so close to zero that for practical purposes it is indistinguishable from zero. Let’s consider a noisy disk drive that transmits each bit correctly with probability (1−f ) and incorrectly with probability f . This model communication channel is known as the binary symmetric channel (figure 1.4).

x

0 -0 @ 

R @ 1 -1

y

P (y = 0 | x = 0) = 1 − f ; P (y = 0 | x = 1) = f ; P (y = 1 | x = 0) = f ; P (y = 1 | x = 1) = 1 − f.

0

(1 − f ) @

1

-0 

@f @ R @ -1

(1 − f )

As an example, let’s imagine that f = 0.1, that is, ten per cent of the bits are flipped (figure 1.5). A useful disk drive would flip no bits at all in its entire lifetime. If we expect to read and write a gigabyte per day for ten years, we require a bit error probability of the order of 10 −15 , or smaller. There are two approaches to this goal.

The physical solution The physical solution is to improve the physical characteristics of the communication channel to reduce its error probability. We could improve our disk drive by 1. using more reliable components in its circuitry; 2. evacuating the air from the disk enclosure so as to eliminate the turbulence that perturbs the reading head from the track; 3. using a larger magnetic patch to represent each bit; or 4. using higher-power signals or cooling the circuitry in order to reduce thermal noise. These physical modifications typically increase the cost of the communication channel.

The ‘system’ solution Information theory and coding theory offer an alternative (and much more exciting) approach: we accept the given noisy channel as it is and add communication systems to it so that we can detect and correct the errors introduced by the channel. As shown in figure 1.6, we add an encoder before the channel and a decoder after it. The encoder encodes the source message s into a transmitted message t, adding redundancy to the original message in some way. The channel adds noise to the transmitted message, yielding a received message r. The decoder uses the known redundancy introduced by the encoding system to infer both the original signal s and the added noise.

Figure 1.4. The binary symmetric channel. The transmitted symbol is x and the received symbol y. The noise level, the probability that a bit is flipped, is f . Figure 1.5. A binary data sequence of length 10 000 transmitted over a binary symmetric channel with noise level f = 0.1. [Dilbert image c Copyright 1997 United Feature Syndicate, Inc., used with permission.]

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1.2: Error-correcting codes for the binary symmetric channel

Figure 1.6. The ‘system’ solution for achieving reliable communication over a noisy channel. The encoding system introduces systematic redundancy into the transmitted vector t. The decoding system uses this known redundancy to deduce from the received vector r both the original source vector and the noise introduced by the channel.

Source 6

s

ˆs

?

Encoder

t

Decoder 6

r

Noisy channel

-

5

Whereas physical solutions give incremental channel improvements only at an ever-increasing cost, system solutions can turn noisy channels into reliable communication channels with the only cost being a computational requirement at the encoder and decoder. Information theory is concerned with the theoretical limitations and potentials of such systems. ‘What is the best error-correcting performance we could achieve?’ Coding theory is concerned with the creation of practical encoding and decoding systems.

1.2 Error-correcting codes for the binary symmetric channel We now consider examples of encoding and decoding systems. What is the simplest way to add useful redundancy to a transmission? [To make the rules of the game clear: we want to be able to detect and correct errors; and retransmission is not an option. We get only one chance to encode, transmit, and decode.]

Repetition codes A straightforward idea is to repeat every bit of the message a prearranged number of times – for example, three times, as shown in table 1.7. We call this repetition code ‘R3 ’. Imagine that we transmit the source message s=0 0 1 0 1 1 0

Source sequence s

Transmitted sequence t

0 1

000 111

Table 1.7. The repetition code R3 .

over a binary symmetric channel with noise level f = 0.1 using this repetition code. We can describe the channel as ‘adding’ a sparse noise vector n to the transmitted vector – adding in modulo 2 arithmetic, i.e., the binary algebra in which 1+1=0. A possible noise vector n and received vector r = t + n are shown in figure 1.8. s

0 z}|{ t 000 n 000 r 000

0 z}|{ 000 001 001

1 z}|{ 111 000 111

0 z}|{ 000 000 000

1 z}|{ 111 101 010

1 z}|{ 111 000 111

0 z}|{ 000 000 000

How should we decode this received vector? The optimal algorithm looks at the received bits three at a time and takes a majority vote (algorithm 1.9).

Figure 1.8. An example transmission using R3 .

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6

1 — Introduction to Information Theory

Received sequence r

Likelihood ratio γ −3 γ −1 γ −1 γ −1 γ1 γ1 γ1 γ3

000 001 010 100 101 110 011 111

P (r | s = 1) P (r | s = 0)

Decoded sequence ˆs 0 0 0 0 1 1 1 1

At the risk of explaining the obvious, let’s prove this result. The optimal decoding decision (optimal in the sense of having the smallest probability of being wrong) is to find which value of s is most probable, given r. Consider the decoding of a single bit s, which was encoded as t(s) and gave rise to three received bits r = r1 r2 r3 . By Bayes’ theorem, the posterior probability of s is P (s | r1 r2 r3 ) =

P (r1 r2 r3 | s)P (s) . P (r1 r2 r3 )

(1.18)

We can spell out the posterior probability of the two alternatives thus: P (s = 1 | r1 r2 r3 ) =

P (r1 r2 r3 | s = 1)P (s = 1) ; P (r1 r2 r3 )

(1.19)

P (s = 0 | r1 r2 r3 ) =

P (r1 r2 r3 | s = 0)P (s = 0) . P (r1 r2 r3 )

(1.20)

This posterior probability is determined by two factors: the prior probability P (s), and the data-dependent term P (r1 r2 r3 | s), which is called the likelihood of s. The normalizing constant P (r1 r2 r3 ) needn’t be computed when finding the optimal decoding decision, which is to guess sˆ = 0 if P (s = 0 | r) > P (s = 1 | r), and sˆ = 1 otherwise. To find P (s = 0 | r) and P (s = 1 | r), we must make an assumption about the prior probabilities of the two hypotheses s = 0 and s = 1, and we must make an assumption about the probability of r given s. We assume that the prior probabilities are equal: P (s = 0) = P (s = 1) = 0.5; then maximizing the posterior probability P (s | r) is equivalent to maximizing the likelihood P (r | s). And we assume that the channel is a binary symmetric channel with noise level f < 0.5, so that the likelihood is P (r | s) = P (r | t(s)) =

N Y

n=1

P (rn | tn (s)),

(1.21)

where N = 3 is the number of transmitted bits in the block we are considering, and  (1−f ) if rn = tn (1.22) P (rn | tn ) = f if rn 6= tn . Thus the likelihood ratio for the two hypotheses is N Y P (rn | tn (1)) P (r | s = 1) = ; P (r | s = 0) n=1 P (rn | tn (0))

each factor (1−f ) f

P (rn |tn (1)) P (rn |tn (0))

equals

(1−f ) f

if rn = 1 and

f (1−f )

(1.23) if rn = 0. The ratio

is greater than 1, since f < 0.5, so the winning hypothesis is the γ ≡ one with the most ‘votes’, each vote counting for a factor of γ in the likelihood ratio.

Algorithm 1.9. Majority-vote decoding algorithm for R3 . Also shown are the likelihood ratios (1.23), assuming the channel is a binary symmetric channel; γ ≡ (1 − f )/f .

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1.2: Error-correcting codes for the binary symmetric channel

7

Thus the majority-vote decoder shown in algorithm 1.9 is the optimal decoder if we assume that the channel is a binary symmetric channel and that the two possible source messages 0 and 1 have equal prior probability.

We now apply the majority vote decoder to the received vector of figure 1.8. The first three received bits are all 0, so we decode this triplet as a 0. In the second triplet of figure 1.8, there are two 0s and one 1, so we decode this triplet as a 0 – which in this case corrects the error. Not all errors are corrected, however. If we are unlucky and two errors fall in a single block, as in the fifth triplet of figure 1.8, then the decoding rule gets the wrong answer, as shown in figure 1.10. s

0 z}|{ t 000 n 000 r 0 00 |{z}

ˆs corrected errors undetected errors

0

0 z}|{ 000 001 0 01 |{z} 0 ?

1 z}|{ 111 000 1 11 |{z} 1

0 z}|{ 000 000 0 00 |{z} 0

1 z}|{ 111 101 0 10 |{z} 0

?

1 z}|{ 111 000 1 11 |{z} 1

0

Exercise 1.2.[2, p.16] Show that the error probability is reduced by the use of R3 by computing the error probability of this code for a binary symmetric channel with noise level f . The error probability is dominated by the probability that two bits in a block of three are flipped, which scales as f 2 . In the case of the binary symmetric channel with f = 0.1, the R 3 code has a probability of error, after decoding, of pb ' 0.03 per bit. Figure 1.11 shows the result of transmitting a binary image over a binary symmetric channel using the repetition code.

s

encoder -

t

channel f = 10% -

r

Figure 1.10. Decoding the received vector from figure 1.8.

0 z}|{ 000 000 0 00 |{z}

decoder

The exercise’s rating, e.g.‘[2 ]’, indicates its difficulty: ‘1’ exercises are the easiest. Exercises that are accompanied by a marginal rat are especially recommended. If a solution or partial solution is provided, the page is indicated after the difficulty rating; for example, this exercise’s solution is on page 16.

ˆs

-

Figure 1.11. Transmitting 10 000 source bits over a binary symmetric channel with f = 10% using a repetition code and the majority vote decoding algorithm. The probability of decoded bit error has fallen to about 3%; the rate has fallen to 1/3.

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8

1 — Introduction to Information Theory

0.1

0.1 0.01

R1

0.08

R5

1e-05

R1 R3

more useful codes

pb 0.06

0.04

1e-10 R3

0.02 R5 R61

more useful codes

0

R61

1e-15 0

0.2

0.4

0.6 Rate

0.8

1

0

0.2

0.4

0.6 Rate

0.8

1

The repetition code R3 has therefore reduced the probability of error, as desired. Yet we have lost something: our rate of information transfer has fallen by a factor of three. So if we use a repetition code to communicate data over a telephone line, it will reduce the error frequency, but it will also reduce our communication rate. We will have to pay three times as much for each phone call. Similarly, we would need three of the original noisy gigabyte disk drives in order to create a one-gigabyte disk drive with p b = 0.03. Can we push the error probability lower, to the values required for a sellable disk drive – 10−15 ? We could achieve lower error probabilities by using repetition codes with more repetitions. Exercise 1.3.[3, p.16] (a) Show that the probability of error of R N , the repetition code with N repetitions, is pb =

N X

n=(N +1)/2

  N n f (1 − f )N −n , n

(1.24)

for odd N . (b) Assuming f = 0.1, which of the terms in this sum is the biggest? How much bigger is it than the second-biggest term?  (c) Use Stirling’s approximation (p.2) to approximate the N n in the largest term, and find, approximately, the probability of error of the repetition code with N repetitions. (d) Assuming f = 0.1, find how many repetitions are required to get the probability of error down to 10−15 . [Answer: about 60.] So to build a single gigabyte disk drive with the required reliability from noisy gigabyte drives with f = 0.1, we would need sixty of the noisy disk drives. The tradeoff between error probability and rate for repetition codes is shown in figure 1.12.

Block codes – the (7, 4) Hamming code We would like to communicate with tiny probability of error and at a substantial rate. Can we improve on repetition codes? What if we add redundancy to blocks of data instead of encoding one bit at a time? We now study a simple block code.

Figure 1.12. Error probability pb versus rate for repetition codes over a binary symmetric channel with f = 0.1. The right-hand figure shows pb on a logarithmic scale. We would like the rate to be large and pb to be small.

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1.2: Error-correcting codes for the binary symmetric channel

9

A block code is a rule for converting a sequence of source bits s, of length K, say, into a transmitted sequence t of length N bits. To add redundancy, we make N greater than K. In a linear block code, the extra N − K bits are linear functions of the original K bits; these extra bits are called parity-check bits. An example of a linear block code is the (7, 4) Hamming code, which transmits N = 7 bits for every K = 4 source bits. t5 s1 t7

Figure 1.13. Pictorial representation of encoding for the (7, 4) Hamming code.

1

s3 s4

s2

1

0 0

t

1

6

0

0

(b)

(a)

The encoding operation for the code is shown pictorially in figure 1.13. We arrange the seven transmitted bits in three intersecting circles. The first four transmitted bits, t1 t2 t3 t4 , are set equal to the four source bits, s 1 s2 s3 s4 . The parity-check bits t5 t6 t7 are set so that the parity within each circle is even: the first parity-check bit is the parity of the first three source bits (that is, it is 0 if the sum of those bits is even, and 1 if the sum is odd); the second is the parity of the last three; and the third parity bit is the parity of source bits one, three and four. As an example, figure 1.13b shows the transmitted codeword for the case s = 1000. Table 1.14 shows the codewords generated by each of the 2 4 = sixteen settings of the four source bits. These codewords have the special property that any pair differ from each other in at least three bits. s

t

s

t

s

t

s

t

0000 0001 0010 0011

0000000 0001011 0010111 0011100

0100 0101 0110 0111

0100110 0101101 0110001 0111010

1000 1001 1010 1011

1000101 1001110 1010010 1011001

1100 1101 1110 1111

1100011 1101000 1110100 1111111

Because the Hamming code is a linear code, it can be written compactly in terms of matrices as follows. The transmitted codeword t is obtained from the source sequence s by a linear operation, t = GTs, where G is the generator matrix of the code,  1 0 0  0 1 0   0 0 1  T G =  0 0 0  1 1 1   0 1 1 1 0 1

(1.25)

0 0 0 1 0 1 1



    ,    

(1.26)

and the encoding operation (1.25) uses modulo-2 arithmetic (1 + 1 = 0, 0 + 1 = 1, etc.). In the encoding operation (1.25) I have assumed that s and t are column vectors. If instead they are row vectors, then this equation is replaced by t = sG,

(1.27)

Table 1.14. The sixteen codewords {t} of the (7, 4) Hamming code. Any pair of codewords differ from each other in at least three bits.

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10

1 — Introduction to Information Theory where



1  0  G= 0 0

0 1 0 0

0 0 1 0

0 0 0 1

1 1 1 0

0 1 1 1

 1 0  . 1  1

(1.28)

I find it easier to relate to the right-multiplication (1.25) than the left-multiplication (1.27). Many coding theory texts use the left-multiplying conventions (1.27–1.28), however. The rows of the generator matrix (1.28) can be viewed as defining four basis vectors lying in a seven-dimensional binary space. The sixteen codewords are obtained by making all possible linear combinations of these vectors.

Decoding the (7, 4) Hamming code When we invent a more complex encoder s → t, the task of decoding the received vector r becomes less straightforward. Remember that any of the bits may have been flipped, including the parity bits. If we assume that the channel is a binary symmetric channel and that all source vectors are equiprobable, then the optimal decoder identifies the source vector s whose encoding t(s) differs from the received vector r in the fewest bits. [Refer to the likelihood function (1.23) to see why this is so.] We could solve the decoding problem by measuring how far r is from each of the sixteen codewords in table 1.14, then picking the closest. Is there a more efficient way of finding the most probable source vector? Syndrome decoding for the Hamming code For the (7, 4) Hamming code there is a pictorial solution to the decoding problem, based on the encoding picture, figure 1.13. As a first example, let’s assume the transmission was t = 1000101 and the noise flips the second bit, so the received vector is r = 1000101 ⊕ 0100000 = 1100101. We write the received vector into the three circles as shown in figure 1.15a, and look at each of the three circles to see whether its parity is even. The circles whose parity is not even are shown by dashed lines in figure 1.15b. The decoding task is to find the smallest set of flipped bits that can account for these violations of the parity rules. [The pattern of violations of the parity checks is called the syndrome, and can be written as a binary vector – for example, in figure 1.15b, the syndrome is z = (1, 1, 0), because the first two circles are ‘unhappy’ (parity 1) and the third circle is ‘happy’ (parity 0).] To solve the decoding task, we ask the question: can we find a unique bit that lies inside all the ‘unhappy’ circles and outside all the ‘happy’ circles? If so, the flipping of that bit would account for the observed syndrome. In the case shown in figure 1.15b, the bit r2 lies inside the two unhappy circles and outside the happy circle; no other single bit has this property, so r 2 is the only single bit capable of explaining the syndrome. Let’s work through a couple more examples. Figure 1.15c shows what happens if one of the parity bits, t5 , is flipped by the noise. Just one of the checks is violated. Only r5 lies inside this unhappy circle and outside the other two happy circles, so r5 is identified as the only single bit capable of explaining the syndrome. If the central bit r3 is received flipped, figure 1.15d shows that all three checks are violated; only r3 lies inside all three circles, so r3 is identified as the suspect bit.

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1.2: Error-correcting codes for the binary symmetric channel

11

r5 r1

r3

r2 r

r4

r7

6

(a)

0*

1 1*

1

0

1

0 1

1 1

1*

0 0

0

1

(b)

1

0

0

(c)

0*

1* 0

0

0

(d)

1 1

0

1 0

-

0

1

1*

1

0*

0

0

Figure 1.15. Pictorial representation of decoding of the Hamming (7, 4) code. The received vector is written into the diagram as shown in (a). In (b,c,d,e), the received vector is shown, assuming that the transmitted vector was as in figure 1.13b and the bits labelled by ? were flipped. The violated parity checks are highlighted by dashed circles. One of the seven bits is the most probable suspect to account for each ‘syndrome’, i.e., each pattern of violated and satisfied parity checks. In examples (b), (c), and (d), the most probable suspect is the one bit that was flipped. In example (e), two bits have been flipped, s3 and t7 . The most probable suspect is r2 , marked by a circle in (e0 ), which shows the output of the decoding algorithm.

0

(e)

(e )

Syndrome z

000

001

010

011

100

101

110

111

Unflip this bit

none

r7

r6

r4

r5

r1

r2

r3

If you try flipping any one of the seven bits, you’ll find that a different syndrome is obtained in each case – seven non-zero syndromes, one for each bit. There is only one other syndrome, the all-zero syndrome. So if the channel is a binary symmetric channel with a small noise level f , the optimal decoder unflips at most one bit, depending on the syndrome, as shown in algorithm 1.16. Each syndrome could have been caused by other noise patterns too, but any other noise pattern that has the same syndrome must be less probable because it involves a larger number of noise events. What happens if the noise actually flips more than one bit? Figure 1.15e shows the situation when two bits, r 3 and r7 , are received flipped. The syndrome, 110, makes us suspect the single bit r 2 ; so our optimal decoding algorithm flips this bit, giving a decoded pattern with three errors as shown in figure 1.15e0 . If we use the optimal decoding algorithm, any two-bit error pattern will lead to a decoded seven-bit vector that contains three errors.

General view of decoding for linear codes: syndrome decoding We can also describe the decoding problem for a linear code in terms of matrices. The first four received bits, r1 r2 r3 r4 , purport to be the four source bits; and the received bits r5 r6 r7 purport to be the parities of the source bits, as defined by the generator matrix G. We evaluate the three parity-check bits for the received bits, r1 r2 r3 r4 , and see whether they match the three received bits, r5 r6 r7 . The differences (modulo 2) between these two triplets are called the syndrome of the received vector. If the syndrome is zero – if all three parity checks are happy – then the received vector is a codeword, and the most probable decoding is

Algorithm 1.16. Actions taken by the optimal decoder for the (7, 4) Hamming code, assuming a binary symmetric channel with small noise level f . The syndrome vector z lists whether each parity check is violated (1) or satisfied (0), going through the checks in the order of the bits r5 , r6 , and r7 .

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12

1 — Introduction to Information Theory

s

encoder -

parity bits

t

r

channel f = 10%

decoder

-

-

      

given by reading out its first four bits. If the syndrome is non-zero, then the noise sequence for this block was non-zero, and the syndrome is our pointer to the most probable error pattern. The computation of the syndrome vector is a linear operation. If we define the 3 × 4 matrix P such that the matrix of equation (1.26) is   I4 T , (1.29) G = P where I4 is the 4 × 4 identity matrix, then the where the parity-check matrix H is given by H arithmetic, −1 ≡ 1, so  1 1 1 0   H = P I3 =  0 1 1 1 1 0 1 1

All the codewords t = GTs of the code satisfy   0 Ht =  0  . 0 [1 ]

. Exercise 1.4.

ˆs

syndrome vector   is z = Hr, = −P I3 ; in modulo 2  1 0 0 0 1 0 . 0 0 1

(1.30)

(1.31)

Prove that this is so by evaluating the 3 × 4 matrix HGT.

Since the received vector r is given by r = GTs + n, the syndrome-decoding problem is to find the most probable noise vector n satisfying the equation Hn = z.

(1.32)

A decoding algorithm that solves this problem is called a maximum-likelihood decoder. We will discuss decoding problems like this in later chapters.

Summary of the (7, 4) Hamming code’s properties Every possible received vector of length 7 bits is either a codeword, or it’s one flip away from a codeword. Since there are three parity constraints, each of which might or might not be violated, there are 2 × 2 × 2 = 8 distinct syndromes. They can be divided into seven non-zero syndromes – one for each of the one-bit error patterns – and the all-zero syndrome, corresponding to the zero-noise case. The optimal decoder takes no action if the syndrome is zero, otherwise it uses this mapping of non-zero syndromes onto one-bit error patterns to unflip the suspect bit.

Figure 1.17. Transmitting 10 000 source bits over a binary symmetric channel with f = 10% using a (7, 4) Hamming code. The probability of decoded bit error is about 7%.

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1.2: Error-correcting codes for the binary symmetric channel

13

There is a decoding error if the four decoded bits sˆ1 , sˆ2 , sˆ3 , sˆ4 do not all match the source bits s1 , s2 , s3 , s4 . The probability of block error pB is the probability that one or more of the decoded bits in one block fail to match the corresponding source bits, pB = P (ˆs 6= s). (1.33) The probability of bit error pb is the average probability that a decoded bit fails to match the corresponding source bit, pb =

K 1 X P (ˆ sk 6= sk ). K

(1.34)

k=1

In the case of the Hamming code, a decoding error will occur whenever the noise has flipped more than one bit in a block of seven. The probability of block error is thus the probability that two or more bits are flipped in a block. This probability scales as O(f 2 ), as did the probability of error for the repetition code R3 . But notice that the Hamming code communicates at a greater rate, R = 4/7. Figure 1.17 shows a binary image transmitted over a binary symmetric channel using the (7, 4) Hamming code. About 7% of the decoded bits are in error. Notice that the errors are correlated: often two or three successive decoded bits are flipped. Exercise 1.5.[1 ] This exercise and the next three refer to the (7, 4) Hamming code. Decode the received strings: (a) r = 1101011 (b) r = 0110110 (c) r = 0100111 (d) r = 1111111. Exercise 1.6.[2, p.17] (a) Calculate the probability of block error p B of the (7, 4) Hamming code as a function of the noise level f and show that to leading order it goes as 21f 2 . (b) [3 ] Show that to leading order the probability of bit error p b goes as 9f 2 . Exercise 1.7.[2, p.19] Find some noise vectors that give the all-zero syndrome (that is, noise vectors that leave all the parity checks unviolated). How many such noise vectors are there? . Exercise 1.8.[2 ] I asserted above that a block decoding error will result whenever two or more bits are flipped in a single block. Show that this is indeed so. [In principle, there might be error patterns that, after decoding, led only to the corruption of the parity bits, with no source bits incorrectly decoded.]

Summary of codes’ performances Figure 1.18 shows the performance of repetition codes and the Hamming code. It also shows the performance of a family of linear block codes that are generalizations of Hamming codes, called BCH codes. This figure shows that we can, using linear block codes, achieve better performance than repetition codes; but the asymptotic situation still looks grim.

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14

1 — Introduction to Information Theory 0.1 0.01

R1

0.1

0.08

R5

1e-05

more useful codes

pb

H(7,4)

R1

H(7,4)

BCH(511,76)

0.06

0.04

Figure 1.18. Error probability pb versus rate R for repetition codes, the (7, 4) Hamming code and BCH codes with blocklengths up to 1023 over a binary symmetric channel with f = 0.1. The righthand figure shows pb on a logarithmic scale.

1e-10

BCH(31,16) R3

0.02 R5

BCH(15,7) more useful codes

BCH(1023,101)

0

1e-15 0

0.2

0.4

0.6 Rate

0.8

1

0

0.2

0.4

0.6 Rate

0.8

1

Exercise 1.9.[4, p.19] Design an error-correcting code and a decoding algorithm for it, estimate its probability of error, and add it to figure 1.18. [Don’t worry if you find it difficult to make a code better than the Hamming code, or if you find it difficult to find a good decoder for your code; that’s the point of this exercise.] Exercise 1.10.[3, p.20] A (7, 4) Hamming code can correct any one error; might there be a (14, 8) code that can correct any two errors? Optional extra: Does the answer to this question depend on whether the code is linear or nonlinear? Exercise 1.11.[4, p.21] Design an error-correcting code, other than a repetition code, that can correct any two errors in a block of size N .

1.3 What performance can the best codes achieve? There seems to be a trade-off between the decoded bit-error probability p b (which we would like to reduce) and the rate R (which we would like to keep large). How can this trade-off be characterized? What points in the (R, p b ) plane are achievable? This question was addressed by Claude Shannon in his pioneering paper of 1948, in which he both created the field of information theory and solved most of its fundamental problems. At that time there was a widespread belief that the boundary between achievable and nonachievable points in the (R, p b ) plane was a curve passing through the origin (R, pb ) = (0, 0); if this were so, then, in order to achieve a vanishingly small error probability p b , one would have to reduce the rate correspondingly close to zero. ‘No pain, no gain.’ However, Shannon proved the remarkable result that the boundary between achievable and nonachievable points meets the R axis at a non-zero value R = C, as shown in figure 1.19. For any channel, there exist codes that make it possible to communicate with arbitrarily small probability of error p b at non-zero rates. The first half of this book (Parts I–III) will be devoted to understanding this remarkable result, which is called the noisy-channel coding theorem.

Example: f = 0.1 The maximum rate at which communication is possible with arbitrarily small pb is called the capacity of the channel. The formula for the capacity of a



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1.4: Summary

15 0.1 0.01

R1

0.1

0.08

R1

R5

1e-05

pb

H(7,4) 0.06

0.04

1e-10 R3

achievable

not achievable

0.02 R5

achievable

not achievable

0

1e-15 0

0.2

0.4

C

0.6 Rate

0.8

1

0

0.2

0.4

C

0.6 Rate

binary symmetric channel with noise level f is   1 1 C(f ) = 1 − H2 (f ) = 1 − f log2 + (1 − f ) log 2 ; f 1−f

0.8

1

(1.35)

the channel we were discussing earlier with noise level f = 0.1 has capacity C ' 0.53. Let us consider what this means in terms of noisy disk drives. The repetition code R3 could communicate over this channel with p b = 0.03 at a rate R = 1/3. Thus we know how to build a single gigabyte disk drive with pb = 0.03 from three noisy gigabyte disk drives. We also know how to make a single gigabyte disk drive with pb ' 10−15 from sixty noisy one-gigabyte drives (exercise 1.3, p.8). And now Shannon passes by, notices us juggling with disk drives and codes and says: ‘What performance are you trying to achieve? 10 −15 ? You don’t need sixty disk drives – you can get that performance with just two disk drives (since 1/2 is less than 0.53). And if you want pb = 10−18 or 10−24 or anything, you can get there with two disk drives too!’ [Strictly, the above statements might not be quite right, since, as we shall see, Shannon proved his noisy-channel coding theorem by studying sequences of block codes with ever-increasing blocklengths, and the required blocklength might be bigger than a gigabyte (the size of our disk drive), in which case, Shannon might say ‘well, you can’t do it with those tiny disk drives, but if you had two noisy terabyte drives, you could make a single high-quality terabyte drive from them’.]

1.4 Summary The (7, 4) Hamming Code By including three parity-check bits in a block of 7 bits it is possible to detect and correct any single bit error in each block.

Shannon’s noisy-channel coding theorem Information can be communicated over a noisy channel at a non-zero rate with arbitrarily small error probability.

Figure 1.19. Shannon’s noisy-channel coding theorem. The solid curve shows the Shannon limit on achievable values of (R, pb ) for the binary symmetric channel with f = 0.1. Rates up to R = C are achievable with arbitrarily small pb . The points show the performance of some textbook codes, as in figure 1.18. The equation defining the Shannon limit (the solid curve) is R = C/(1 − H2 (pb )), where C and H2 are defined in equation (1.35).

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16

1 — Introduction to Information Theory

Information theory addresses both the limitations and the possibilities of communication. The noisy-channel coding theorem, which we will prove in Chapter 10, asserts both that reliable communication at any rate beyond the capacity is impossible, and that reliable communication at all rates up to capacity is possible. The next few chapters lay the foundations for this result by discussing how to measure information content and the intimately related topic of data compression.

1.5 Further exercises . Exercise 1.12.[2, p.21] Consider the repetition code R9 . One way of viewing this code is as a concatenation of R3 with R3 . We first encode the source stream with R3 , then encode the resulting output with R 3 . We could call this code ‘R23 ’. This idea motivates an alternative decoding algorithm, in which we decode the bits three at a time using the decoder for R3 ; then decode the decoded bits from that first decoder using the decoder for R3 . Evaluate the probability of error for this decoder and compare it with the probability of error for the optimal decoder for R 9 . Do the concatenated encoder and decoder for R 23 have advantages over those for R9 ?

1.6 Solutions Solution to exercise 1.2 (p.7). An error is made by R 3 if two or more bits are flipped in a block of three. So the error probability of R 3 is a sum of two terms: the probability that all three bits are flipped, f 3 ; and the probability that exactly two bits are flipped, 3f 2 (1 − f ). [If these expressions are not obvious, see example 1.1 (p.1): the expressions are P (r = 3 | f, N = 3) and P (r = 2 | f, N = 3).] pb = pB = 3f 2 (1 − f ) + f 3 = 3f 2 − 2f 3 .

(1.36)

This probability is dominated for small f by the term 3f 2 . See exercise 2.38 (p.39) for further discussion of this problem. Solution to exercise 1.3 (p.8). The probability of error for the repetition code RN is dominated by the probability that dN/2e bits are flipped, which goes (for odd N ) as   N f (N +1)/2 (1 − f )(N −1)/2 . (1.37) dN/2e  N can be approximated using the binary entropy function: The term K     1 N N ' 2N H2 (K/N ) , (1.38) ≤ 2N H2 (K/N ) ⇒ 2N H2 (K/N ) ≤ K K N +1 √ where this approximation introduces an error of order N – as shown in equation (1.17). So pb = pB ' 2N (f (1 − f ))N/2 = (4f (1 − f ))N/2 .

(1.39) log 10−15

Setting this equal to the required value of 10 −15 we find N ' 2 log 4f (1−f ) = 68. This answer is a little out because the approximation we used overestimated N and we did not distinguish between dN/2e and N/2. K

  Notation: N/2 denotes the smallest integer greater than or equal to N/2.

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1.6: Solutions

17

A slightly more careful answer (short of explicit computation) goes as follows.  N to the next order, we find: Taking the approximation for K   1 N ' 2N p . (1.40) N/2 2πN/4

This approximation can be proved from an accurate version of Stirling’s approximation (1.12), or by considering the binomial distribution with p = 1/2 and noting 1=

X N  K

K

2−N ' 2−N



 N/2   X 2 2 N N √ 2πσ, (1.41) e−r /2σ ' 2−N N/2 N/2 r=−N/2

p

where σ = N/4, from which equation (1.40) follows. The distinction between  dN/2e and N/2 is not important in this term since N K has a maximum at K = N/2.

Then the probability of error (for odd N ) is to leading order   N f (N +1)/2 (1 − f )(N −1)/2 (1.42) pb ' (N +1)/2 1 1 ' 2N p f [f (1 − f )](N −1)/2 ' p f [4f (1 − f )](N −1)/2 . (1.43) πN/2 πN/8

The equation pb = 10−15 can be written

log 10−15 + log

(N − 1)/2 '



πN/8 f

log 4f (1 − f )

(1.44)

ˆ1 = 68: which may be solved for N iteratively, the first iteration starting from N ˆ2 − 1)/2 ' (N

−15 + 1.7 ˆ2 ' 60.9. = 29.9 ⇒ N −0.44

(1.45)

This answer is found to be stable, so N ' 61 is the blocklength at which pb ' 10−15 .

Solution to exercise 1.6 (p.13). (a) The probability of block error of the Hamming code is a sum of six terms – the probabilities that 2, 3, 4, 5, 6, or 7 errors occur in one block. pB =

7   X 7 r=2

r

f r (1 − f )7−r .

(1.46)

To leading order, this goes as pB '

  7 2 f = 21f 2 . 2

(1.47)

(b) The probability of bit error of the Hamming code is smaller than the probability of block error because a block error rarely corrupts all bits in the decoded block. The leading-order behaviour is found by considering the outcome in the most probable case where the noise vector has weight two. The decoder will erroneously flip a third bit, so that the modified received vector (of length 7) differs in three bits from the transmitted vector. That means, if we average over all seven bits, the probability that a randomly chosen bit is flipped is 3/7 times the block error probability, to leading order. Now, what we really care about is the probability that

In equation (1.44), the logarithms can be taken to any base, as long as it’s the same base throughout. In equation (1.45), I use base 10.

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18

1 — Introduction to Information Theory a source bit is flipped. Are parity bits or source bits more likely to be among these three flipped bits, or are all seven bits equally likely to be corrupted when the noise vector has weight two? The Hamming code is in fact completely symmetric in the protection it affords to the seven bits (assuming a binary symmetric channel). [This symmetry can be proved by showing that the role of a parity bit can be exchanged with a source bit and the resulting code is still a (7, 4) Hamming code; see below.] The probability that any one bit ends up corrupted is the same for all seven bits. So the probability of bit error (for the source bits) is simply three sevenths of the probability of block error. 3 pb ' pB ' 9f 2 . 7

(1.48)

Symmetry of the Hamming (7, 4) code To prove that the (7, 4) code protects all bits equally, we start from the paritycheck matrix   1 1 1 0 1 0 0 H =  0 1 1 1 0 1 0 . (1.49) 1 0 1 1 0 0 1

The symmetry among the seven transmitted bits will be easiest to see if we reorder the seven bits using the permutation (t 1 t2 t3 t4 t5 t6 t7 ) → (t5 t2 t3 t4 t1 t6 t7 ). Then we can rewrite H thus:   1 1 1 0 1 0 0 (1.50) H =  0 1 1 1 0 1 0 . 0 0 1 1 1 0 1

Now, if we take any two parity constraints that t satisfies and add them together, we get another parity constraint. For example, row 1 asserts t 5 + t2 + t3 + t1 = even, and row 2 asserts t2 + t3 + t4 + t6 = even, and the sum of these two constraints is t5 + 2t2 + 2t3 + t1 + t4 + t6 = even;

(1.51)

we can drop the terms 2t2 and 2t3 , since they are even whatever t2 and t3 are; thus we have derived the parity constraint t 5 + t1 + t4 + t6 = even, which we can if we wish add into the parity-check matrix as a fourth row. [The set of vectors satisfying Ht = 0 will not be changed.] We thus define   1 1 1 0 1 0 0  0 1 1 1 0 1 0   H0 =  (1.52)  0 0 1 1 1 0 1 . 1 0 0 1 1 1 0 The fourth row is the sum (modulo two) of the top two rows. Notice that the second, third, and fourth rows are all cyclic shifts of the top row. If, having added the fourth redundant constraint, we drop the first constraint, we obtain a new parity-check matrix H00 ,   0 1 1 1 0 1 0 H00 =  0 0 1 1 1 0 1  , (1.53) 1 0 0 1 1 1 0 which still satisfies H00 t = 0 for all codewords, and which looks just like the starting H in (1.50), except that all the columns have shifted along one

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1.6: Solutions

19

to the right, and the rightmost column has reappeared at the left (a cyclic permutation of the columns). This establishes the symmetry among the seven bits. Iterating the above procedure five more times, we can make a total of seven different H matrices for the same original code, each of which assigns each bit to a different role. We may also construct the super-redundant seven-row parity-check matrix for the code,   1 1 1 0 1 0 0  0 1 1 1 0 1 0     0 0 1 1 1 0 1     (1.54) H000 =   1 0 0 1 1 1 0 .  0 1 0 0 1 1 1     1 0 1 0 0 1 1  1 1 0 1 0 0 1 This matrix is ‘redundant’ in the sense that the space spanned by its rows is only three-dimensional, not seven. This matrix is also a cyclic matrix. Every row is a cyclic permutation of the top row.

Cyclic codes: if there is an ordering of the bits t 1 . . . tN such that a linear code has a cyclic parity-check matrix, then the code is called a cyclic code. The codewords of such a code also have cyclic properties: any cyclic permutation of a codeword is a codeword. For example, the Hamming (7, 4) code, with its bits ordered as above, consists of all seven cyclic shifts of the codewords 1110100 and 1011000, and the codewords 0000000 and 1111111. Cyclic codes are a cornerstone of the algebraic approach to error-correcting codes. We won’t use them again in this book, however, as they have been superceded by sparse-graph codes (Part VI). Solution to exercise 1.7 (p.13). There are fifteen non-zero noise vectors which give the all-zero syndrome; these are precisely the fifteen non-zero codewords of the Hamming code. Notice that because the Hamming code is linear , the sum of any two codewords is a codeword.

Graphs corresponding to codes Solution to exercise 1.9 (p.14). When answering this question, you will probably find that it is easier to invent new codes than to find optimal decoders for them. There are many ways to design codes, and what follows is just one possible train of thought. We make a linear block code that is similar to the (7, 4) Hamming code, but bigger. Many codes can be conveniently expressed in terms of graphs. In figure 1.13, we introduced a pictorial representation of the (7, 4) Hamming code. If we replace that figure’s big circles, each of which shows that the parity of four particular bits is even, by a ‘parity-check node’ that is connected to the four bits, then we obtain the representation of the (7, 4) Hamming code by a bipartite graph as shown in figure 1.20. The 7 circles are the 7 transmitted bits. The 3 squares are the parity-check nodes (not to be confused with the 3 parity-check bits, which are the three most peripheral circles). The graph is a ‘bipartite’ graph because its nodes fall into two classes – bits and checks

Figure 1.20. The graph of the (7, 4) Hamming code. The 7 circles are the bit nodes and the 3 squares are the parity-check nodes.

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20

1 — Introduction to Information Theory

– and there are edges only between nodes in different classes. The graph and the code’s parity-check matrix (1.30) are simply related to each other: each parity-check node corresponds to a row of H and each bit node corresponds to a column of H; for every 1 in H, there is an edge between the corresponding pair of nodes. Having noticed this connection between linear codes and graphs, one way to invent linear codes is simply to think of a bipartite graph. For example, a pretty bipartite graph can be obtained from a dodecahedron by calling the vertices of the dodecahedron the parity-check nodes, and putting a transmitted bit on each edge in the dodecahedron. This construction defines a paritycheck matrix in which every column has weight 2 and every row has weight 3. [The weight of a binary vector is the number of 1s it contains.] This code has N = 30 bits, and it appears to have M apparent = 20 paritycheck constraints. Actually, there are only M = 19 independent constraints; the 20th constraint is redundant (that is, if 19 constraints are satisfied, then the 20th is automatically satisfied); so the number of source bits is K = N − M = 11. The code is a (30, 11) code. It is hard to find a decoding algorithm for this code, but we can estimate its probability of error by finding its lowest-weight codewords. If we flip all the bits surrounding one face of the original dodecahedron, then all the parity checks will be satisfied; so the code has 12 codewords of weight 5, one for each face. Since the lowest-weight codewords have weight 5, we say that the code has distance d = 5; the (7, 4) Hamming code had distance 3 and could correct all single bit-flip errors. A code with distance 5 can correct all double bit-flip errors, but there are some triple bit-flip errors that it cannot correct. So the error probability of this code, assuming a binary symmetric channel, will be dominated, at least for low noise levels f , by a term of order f 3 , perhaps something like   5 3 12 f (1 − f )27 . (1.55) 3 Of course, there is no obligation to make codes whose graphs can be represented on a plane, as this one can; the best linear codes, which have simple graphical descriptions, have graphs that are more tangled, as illustrated by the tiny (16, 4) code of figure 1.22. Furthermore, there is no reason for sticking to linear codes; indeed some nonlinear codes – codes whose codewords cannot be defined by a linear equation like Ht = 0 – have very good properties. But the encoding and decoding of a nonlinear code are even trickier tasks. Solution to exercise 1.10 (p.14). First let’s assume we are making a linear code and decoding it with syndrome decoding. If there are N transmitted bits, then the number of possible error patterns of weight up to two is       N N N . (1.56) + + 0 1 2 For N = 14, that’s 91 + 14 + 1 = 106 patterns. Now, every distinguishable error pattern must give rise to a distinct syndrome; and the syndrome is a list of M bits, so the maximum possible number of syndromes is 2 M . For a (14, 8) code, M = 6, so there are at most 2 6 = 64 syndromes. The number of possible error patterns of weight up to two, 106, is bigger than the number of syndromes, 64, so we can immediately rule out the possibility that there is a (14, 8) code that is 2-error-correcting.

Figure 1.21. The graph defining the (30, 11) dodecahedron code. The circles are the 30 transmitted bits and the triangles are the 20 parity checks. One parity check is redundant.

Figure 1.22. Graph of a rate-1/4 low-density parity-check code (Gallager code) with blocklength N = 16, and M = 12 parity-check constraints. Each white circle represents a transmitted bit. Each bit participates in j = 3 constraints, represented by squares. The edges between nodes were placed at random. (See Chapter 47 for more.)

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1.6: Solutions

21

The same counting argument works fine for nonlinear codes too. When the decoder receives r = t + n, his aim is to deduce both t and n from r. If it is the case that the sender can select any transmission t from a code of size St , and the channel can select any noise vector from a set of size S n , and those two selections can be recovered from the received bit string r, which is one of at most 2N possible strings, then it must be the case that St Sn ≤ 2 N .

(1.57)

So, for a (N, K) two-error-correcting code, whether linear or nonlinear,       N N N ≤ 2N . (1.58) 2K + + 0 1 2 Solution to exercise 1.11 (p.14). There are various strategies for making codes that can correct multiple errors, and I strongly recommend you think out one or two of them for yourself. If your approach uses a linear code, e.g., one with a collection of M parity checks, it is helpful to bear in mind the counting argument given in the previous exercise, in order to anticipate how many parity checks, M , you might need. Examples of codes that can correct any two errors are the (30, 11) dodecahedron code on page 20, and the (15, 6) pentagonful code to be introduced on p.221. Further simple ideas for making codes that can correct multiple errors from codes that can correct only one error are discussed in section 13.7. Solution to exercise 1.12 (p.16). The probability of error of R 23 is, to leading order, (1.59) pb (R23 ) ' 3 [pb (R3 )]2 = 3(3f 2 )2 + · · · = 27f 4 + · · · , whereas the probability of error of R 9 is dominated by the probability of five flips,   9 5 pb (R9 ) ' f (1 − f )4 ' 126f 5 + · · · . (1.60) 5 The R23 decoding procedure is therefore suboptimal, since there are noise vectors of weight four that cause it to make a decoding error. It has the advantage, however, of requiring smaller computational resources: only memorization of three bits, and counting up to three, rather than counting up to nine. This simple code illustrates an important concept. Concatenated codes are widely used in practice because concatenation allows large codes to be implemented using simple encoding and decoding hardware. Some of the best known practical codes are concatenated codes.

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2 Probability, Entropy, and Inference This chapter, and its sibling, Chapter 8, devote some time to notation. Just as the White Knight distinguished between the song, the name of the song, and what the name of the song was called (Carroll, 1998), we will sometimes need to be careful to distinguish between a random variable, the value of the random variable, and the proposition that asserts that the random variable has a particular value. In any particular chapter, however, I will use the most simple and friendly notation possible, at the risk of upsetting pure-minded readers. For example, if something is ‘true with probability 1’, I will usually simply say that it is ‘true’.

2.1 Probabilities and ensembles An ensemble X is a triple (x, AX , PX ), where the outcome x is the value of a random variable, which takes on one of a set of possible values, AX = {a1 , a2 , . . . , ai , . . . , aI }, having P probabilities PX = {p1 , p2 , . . . , pI }, with P (x = ai ) = pi , pi ≥ 0 and ai ∈AX P (x = ai ) = 1. The name A is mnemonic for ‘alphabet’. One example of an ensemble is a letter that is randomly selected from an English document. This ensemble is shown in figure 2.1. There are twenty-seven possible letters: a–z, and a space character ‘-’.

Abbreviations. Briefer notation will sometimes be used. P (x = ai ) may be written as P (ai ) or P (x).

For example,

Probability of a subset. If T is a subset of A X then: X P (T ) = P (x ∈ T ) = P (x = ai ).

(2.1)

ai ∈T

For example, if we define V to be vowels from figure 2.1, V {a, e, i, o, u}, then P (V ) = 0.06 + 0.09 + 0.06 + 0.07 + 0.03 = 0.31.

(2.2)

We call P (x, y) the joint probability of x and y.

N.B. In a joint ensemble XY the two variables are not necessarily independent. 22

ai

pi

1 2 3 4 5 6 7 8 9 10 11 12 13 14 15 16 17 18 19 20 21 22 23 24 25 26 27

a b c d e f g h i j k l m n o p q r s t u v w x y z –

0.0575 0.0128 0.0263 0.0285 0.0913 0.0173 0.0133 0.0313 0.0599 0.0006 0.0084 0.0335 0.0235 0.0596 0.0689 0.0192 0.0008 0.0508 0.0567 0.0706 0.0334 0.0069 0.0119 0.0073 0.0164 0.0007 0.1928

a b c d e f g h i j k l m n o p q r s t u v w x y z –

=

A joint ensemble XY is an ensemble in which each outcome is an ordered pair x, y with x ∈ AX = {a1 , . . . , aI } and y ∈ AY = {b1 , . . . , bJ }. Commas are optional when writing ordered pairs, so xy ⇔ x, y.

i

Figure 2.1. Probability distribution over the 27 outcomes for a randomly selected letter in an English language document (estimated from The Frequently Asked Questions Manual for Linux ). The picture shows the probabilities by the areas of white squares.

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2.1: Probabilities and ensembles

23

x

Figure 2.2. The probability distribution over the 27×27 possible bigrams xy in an English language document, The Frequently Asked Questions Manual for Linux.

a b c d e f g h i j k l m n o p q r s t u v w x y z – abcdefghijklmnopqrstuvwxyz– y

Marginal probability. We can obtain the marginal probability P (x) from the joint probability P (x, y) by summation: X P (x = ai ) ≡ P (x = ai , y). (2.3) y∈AY

Similarly, using briefer notation, the marginal probability of y is: X P (y) ≡ P (x, y). (2.4) x∈AX

Conditional probability P (x = ai | y = bj ) ≡

P (x = ai , y = bj ) if P (y = bj ) 6= 0. P (y = bj )

(2.5)

[If P (y = bj ) = 0 then P (x = ai | y = bj ) is undefined.]

We pronounce P (x = ai | y = bj ) ‘the probability that x equals ai , given y equals bj ’.

Example 2.1. An example of a joint ensemble is the ordered pair XY consisting of two successive letters in an English document. The possible outcomes are ordered pairs such as aa, ab, ac, and zz; of these, we might expect ab and ac to be more probable than aa and zz. An estimate of the joint probability distribution for two neighbouring characters is shown graphically in figure 2.2. This joint ensemble has the special property that its two marginal distributions, P (x) and P (y), are identical. They are both equal to the monogram distribution shown in figure 2.1. From this joint ensemble P (x, y) we can obtain conditional distributions, P (y | x) and P (x | y), by normalizing the rows and columns, respectively (figure 2.3). The probability P (y | x = q) is the probability distribution of the second letter given that the first letter is a q. As you can see in figure 2.3a, the two most probable values for the second letter y given

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24

2 — Probability, Entropy, and Inference

x a b c d e f g h i j k l m n o p q r s t u v w x y z –

x a b c d e f g h i j k l m n o p q r s t u v w x y z –

Figure 2.3. Conditional probability distributions. (a) P (y | x): Each row shows the conditional distribution of the second letter, y, given the first letter, x, in a bigram xy. (b) P (x | y): Each column shows the conditional distribution of the first letter, x, given the second letter, y.

abcdefghijklmnopqrstuvwxyz– y

abcdefghijklmnopqrstuvwxyz– y

(a) P (y | x)

(b) P (x | y)

that the first letter x is q are u and -. (The space is common after q because the source document makes heavy use of the word FAQ.) The probability P (x | y = u) is the probability distribution of the first letter x given that the second letter y is a u. As you can see in figure 2.3b the two most probable values for x given y = u are n and o. Rather than writing down the joint probability directly, we often define an ensemble in terms of a collection of conditional probabilities. The following rules of probability theory will be useful. (H denotes assumptions on which the probabilities are based.) Product rule – obtained from the definition of conditional probability: P (x, y | H) = P (x | y, H)P (y | H) = P (y | x, H)P (x | H).

(2.6)

This rule is also known as the chain rule. Sum rule – a rewriting of the marginal probability definition: X P (x | H) = P (x, y | H)

(2.7)

y

=

X y

P (x | y, H)P (y | H).

(2.8)

Bayes’ theorem – obtained from the product rule: P (y | x, H) = =

P (x | y, H)P (y | H) P (x | H) P (x | y, H)P (y | H) P . 0 0 y 0 P (x | y , H)P (y | H)

(2.9) (2.10)

Independence. Two random variables X and Y are independent (sometimes written X⊥Y ) if and only if P (x, y) = P (x)P (y).

(2.11)

Exercise 2.2.[1, p.40] Are the random variables X and Y in the joint ensemble of figure 2.2 independent?

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2.2: The meaning of probability

25

I said that we often define an ensemble in terms of a collection of conditional probabilities. The following example illustrates this idea. Example 2.3. Jo has a test for a nasty disease. We denote Jo’s state of health by the variable a and the test result by b. a=1 a=0

Jo has the disease Jo does not have the disease.

(2.12)

The result of the test is either ‘positive’ (b = 1) or ‘negative’ (b = 0); the test is 95% reliable: in 95% of cases of people who really have the disease, a positive result is returned, and in 95% of cases of people who do not have the disease, a negative result is obtained. The final piece of background information is that 1% of people of Jo’s age and background have the disease. OK – Jo has the test, and the result is positive. What is the probability that Jo has the disease? Solution. We write down all the provided probabilities. The test reliability specifies the conditional probability of b given a: P (b = 1 | a = 1) = 0.95 P (b = 0 | a = 1) = 0.05

P (b = 1 | a = 0) = 0.05 P (b = 0 | a = 0) = 0.95;

(2.13)

and the disease prevalence tells us about the marginal probability of a: P (a = 1) = 0.01

P (a = 0) = 0.99.

(2.14)

From the marginal P (a) and the conditional probability P (b | a) we can deduce the joint probability P (a, b) = P (a)P (b | a) and any other probabilities we are interested in. For example, by the sum rule, the marginal probability of b = 1 – the probability of getting a positive result – is P (b = 1) = P (b = 1 | a = 1)P (a = 1) + P (b = 1 | a = 0)P (a = 0).

(2.15)

Jo has received a positive result b = 1 and is interested in how plausible it is that she has the disease (i.e., that a = 1). The man in the street might be duped by the statement ‘the test is 95% reliable, so Jo’s positive result implies that there is a 95% chance that Jo has the disease’, but this is incorrect. The correct solution to an inference problem is found using Bayes’ theorem. P (b = 1 | a = 1)P (a = 1) (2.16) P (b = 1 | a = 1)P (a = 1) + P (b = 1 | a = 0)P (a = 0) 0.95 × 0.01 (2.17) = 0.95 × 0.01 + 0.05 × 0.99 = 0.16. (2.18)

P (a = 1 | b = 1) =

So in spite of the positive result, the probability that Jo has the disease is only 16%. 2

2.2 The meaning of probability Probabilities can be used in two ways. Probabilities can describe frequencies of outcomes in random experiments, but giving noncircular definitions of the terms ‘frequency’ and ‘random’ is a challenge – what does it mean to say that the frequency of a tossed coin’s

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26

2 — Probability, Entropy, and Inference

Notation. Let ‘the degree of belief in proposition x’ be denoted by B(x). The negation of x (not-x) is written x. The degree of belief in a conditional proposition, ‘x, assuming proposition y to be true’, is represented by B(x | y). Axiom 1. Degrees of belief can be ordered; if B(x) is ‘greater’ than B(y), and B(y) is ‘greater’ than B(z), then B(x) is ‘greater’ than B(z). [Consequence: beliefs can be mapped onto real numbers.] Axiom 2. The degree of belief in a proposition x and its negation x are related. There is a function f such that B(x) = f [B(x)]. Axiom 3. The degree of belief in a conjunction of propositions x, y (x and y) is related to the degree of belief in the conditional proposition x | y and the degree of belief in the proposition y. There is a function g such that B(x, y) = g [B(x | y), B(y)] .

coming up heads is 1/2? If we say that this frequency is the average fraction of heads in long sequences, we have to define ‘average’; and it is hard to define ‘average’ without using a word synonymous to probability! I will not attempt to cut this philosophical knot. Probabilities can also be used, more generally, to describe degrees of belief in propositions that do not involve random variables – for example ‘the probability that Mr. S. was the murderer of Mrs. S., given the evidence’ (he either was or wasn’t, and it’s the jury’s job to assess how probable it is that he was); ‘the probability that Thomas Jefferson had a child by one of his slaves’; ‘the probability that Shakespeare’s plays were written by Francis Bacon’; or, to pick a modern-day example, ‘the probability that a particular signature on a particular cheque is genuine’. The man in the street is happy to use probabilities in both these ways, but some books on probability restrict probabilities to refer only to frequencies of outcomes in repeatable random experiments. Nevertheless, degrees of belief can be mapped onto probabilities if they satisfy simple consistency rules known as the Cox axioms (Cox, 1946) (figure 2.4). Thus probabilities can be used to describe assumptions, and to describe inferences given those assumptions. The rules of probability ensure that if two people make the same assumptions and receive the same data then they will draw identical conclusions. This more general use of probability to quantify beliefs is known as the Bayesian viewpoint. It is also known as the subjective interpretation of probability, since the probabilities depend on assumptions. Advocates of a Bayesian approach to data modelling and pattern recognition do not view this subjectivity as a defect, since in their view, you cannot do inference without making assumptions. In this book it will from time to time be taken for granted that a Bayesian approach makes sense, but the reader is warned that this is not yet a globally held view – the field of statistics was dominated for most of the 20th century by non-Bayesian methods in which probabilities are allowed to describe only random variables. The big difference between the two approaches is that

Box 2.4. The Cox axioms. If a set of beliefs satisfy these axioms then they can be mapped onto probabilities satisfying P (false) = 0, P (true) = 1, 0 ≤ P (x) ≤ 1, and the rules of probability: P (x) = 1 − P (x), and P (x, y) = P (x | y)P (y).

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2.3: Forward probabilities and inverse probabilities

27

Bayesians also use probabilities to describe inferences.

2.3 Forward probabilities and inverse probabilities Probability calculations often fall into one of two categories: forward probability and inverse probability. Here is an example of a forward probability problem: Exercise 2.4.[2, p.40] An urn contains K balls, of which B are black and W = K − B are white. Fred draws a ball at random from the urn and replaces it, N times. (a) What is the probability distribution of the number of times a black ball is drawn, nB ? (b) What is the expectation of nB ? What is the variance of nB ? What is the standard deviation of nB ? Give numerical answers for the cases N = 5 and N = 400, when B = 2 and K = 10. Forward probability problems involve a generative model that describes a process that is assumed to give rise to some data; the task is to compute the probability distribution or expectation of some quantity that depends on the data. Here is another example of a forward probability problem: Exercise 2.5.[2, p.40] An urn contains K balls, of which B are black and W = K − B are white. We define the fraction f B ≡ B/K. Fred draws N times from the urn, exactly as in exercise 2.4, obtaining n B blacks, and computes the quantity z=

(nB − fB N )2 . N fB (1 − fB )

(2.19)

What is the expectation of z? In the case N = 5 and f B = 1/5, what is the probability distribution of z? What is the probability that z < 1? [Hint: compare z with the quantities computed in the previous exercise.] Like forward probability problems, inverse probability problems involve a generative model of a process, but instead of computing the probability distribution of some quantity produced by the process, we compute the conditional probability of one or more of the unobserved variables in the process, given the observed variables. This invariably requires the use of Bayes’ theorem. Example 2.6. There are eleven urns labelled by u ∈ {0, 1, 2, . . . , 10}, each containing ten balls. Urn u contains u black balls and 10 − u white balls. Fred selects an urn u at random and draws N times with replacement from that urn, obtaining nB blacks and N − nB whites. Fred’s friend, Bill, looks on. If after N = 10 draws n B = 3 blacks have been drawn, what is the probability that the urn Fred is using is urn u, from Bill’s point of view? (Bill doesn’t know the value of u.) Solution. The joint probability distribution of the random variables u and n B can be written P (u, nB | N ) = P (nB | u, N )P (u). (2.20)

From the joint probability of u and n B , we can obtain the conditional distribution of u given nB : P (u | nB , N ) = =

P (u, nB | N ) P (nB | N ) P (nB | u, N )P (u) . P (nB | N )

(2.21) (2.22)

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28

2 — Probability, Entropy, and Inference u

Figure 2.5. Joint probability of u and nB for Bill and Fred’s urn problem, after N = 10 draws.

0 1 2 3 4 5 6 7 8 9 10 0 1 2 3 4 5 6 7 8 9 10 nB 1 for all u. You wrote down the The marginal probability of u is P (u) = 11 probability of nB given u and N , P (nB | u, N ), when you solved exercise 2.4 (p.27). [You are doing the highly recommended exercises, aren’t you?] If we define fu ≡ u/10 then

P (nB | u, N ) =



 N f nB (1 − fu )N −nB . nB u

(2.23)

P (nB | N ) =

u

P (u, nB | N ) =

X u

P (u)P (nB | u, N ).

(2.24)

So the conditional probability of u given n B is P (u | nB , N ) = =

P (u)P (nB | u, N ) P (nB | N )   1 1 N f nB (1 − fu )N −nB . P (nB | N ) 11 nB u

0.2 0.15 0.1 0.05 0 0

What about the denominator, P (nB | N )? This is the marginal probability of nB , which we can obtain using the sum rule: X

0.3 0.25

(2.25) (2.26)

This conditional distribution can be found by normalizing column 3 of figure 2.5 and is shown in figure 2.6. The normalizing constant, the marginal probability of nB , is P (nB = 3 | N = 10) = 0.083. The posterior probability (2.26) is correct for all u, including the end-points u = 0 and u = 10, where fu = 0 and fu = 1 respectively. The posterior probability that u = 0 given nB = 3 is equal to zero, because if Fred were drawing from urn 0 it would be impossible for any black balls to be drawn. The posterior probability that u = 10 is also zero, because there are no white balls in that urn. The other hypotheses u = 1, u = 2, . . . u = 9 all have non-zero posterior probability. 2

Terminology of inverse probability In inverse probability problems it is convenient to give names to the probabilities appearing in Bayes’ theorem. In equation (2.25), we call the marginal probability P (u) the prior probability of u, and P (n B | u, N ) is called the likelihood of u. It is important to note that the terms likelihood and probability are not synonyms. The quantity P (nB | u, N ) is a function of both nB and u. For fixed u, P (nB | u, N ) defines a probability over nB . For fixed nB , P (nB | u, N ) defines the likelihood of u.

1

2

3

4

5 u

6

7

8

9 10

u

P (u | nB = 3, N )

0 1 2 3 4 5 6 7 8 9 10

0 0.063 0.22 0.29 0.24 0.13 0.047 0.0099 0.00086 0.0000096 0

Figure 2.6. Conditional probability of u given nB = 3 and N = 10.

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2.3: Forward probabilities and inverse probabilities

29

Never say ‘the likelihood of the data’. Always say ‘the likelihood of the parameters’. The likelihood function is not a probability distribution. (If you want to mention the data that a likelihood function is associated with, you may say ‘the likelihood of the parameters given the data’.) The conditional probability P (u | n B , N ) is called the posterior probability of u given nB . The normalizing constant P (nB | N ) has no u-dependence so its value is not important if we simply wish to evaluate the relative probabilities of the alternative hypotheses u. However, in most data-modelling problems of any complexity, this quantity becomes important, and it is given various names: P (nB | N ) is known as the evidence or the marginal likelihood. If θ denotes the unknown parameters, D denotes the data, and H denotes the overall hypothesis space, the general equation: P (D | θ, H)P (θ | H) P (D | H)

(2.27)

likelihood × prior . evidence

(2.28)

P (θ | D, H) = is written: posterior =

Inverse probability and prediction Example 2.6 (continued). Assuming again that Bill has observed n B = 3 blacks in N = 10 draws, let Fred draw another ball from the same urn. What is the probability that the next drawn ball is a black? [You should make use of the posterior probabilities in figure 2.6.] Solution. By the sum rule, P (ballN+1 is black | nB , N ) =

X u

P (ballN+1 is black | u, nB , N )P (u | nB , N ).

(2.29) Since the balls are drawn with replacement from the chosen urn, the probability P (ballN+1 is black | u, nB , N ) is just fu = u/10, whatever nB and N are. So X P (ballN+1 is black | nB , N ) = fu P (u | nB , N ). (2.30) u

Using the values of P (u | nB , N ) given in figure 2.6 we obtain P (ballN+1 is black | nB = 3, N = 10) = 0.333.

2

(2.31)

Comment. Notice the difference between this prediction obtained using probability theory, and the widespread practice in statistics of making predictions by first selecting the most plausible hypothesis (which here would be that the urn is urn u = 3) and then making the predictions assuming that hypothesis to be true (which would give a probability of 0.3 that the next ball is black). The correct prediction is the one that takes into account the uncertainty by marginalizing over the possible values of the hypothesis u. Marginalization here leads to slightly more moderate, less extreme predictions.

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30

2 — Probability, Entropy, and Inference

Inference as inverse probability Now consider the following exercise, which has the character of a simple scientific investigation. Example 2.7. Bill tosses a bent coin N times, obtaining a sequence of heads and tails. We assume that the coin has a probability f H of coming up heads; we do not know fH . If nH heads have occurred in N tosses, what is the probability distribution of f H ? (For example, N might be 10, and nH might be 3; or, after a lot more tossing, we might have N = 300 and nH = 29.) What is the probability that the N +1th outcome will be a head, given nH heads in N tosses? Unlike example 2.6 (p.27), this problem has a subjective element. Given a restricted definition of probability that says ‘probabilities are the frequencies of random variables’, this example is different from the eleven-urns example. Whereas the urn u was a random variable, the bias f H of the coin would not normally be called a random variable. It is just a fixed but unknown parameter that we are interested in. Yet don’t the two examples 2.6 and 2.7 seem to have an essential similarity? [Especially when N = 10 and n H = 3!] To solve example 2.7, we have to make an assumption about what the bias of the coin fH might be. This prior probability distribution over f H , P (fH ), corresponds to the prior over u in the eleven-urns problem. In that example, the helpful problem definition specified P (u). In real life, we have to make assumptions in order to assign priors; these assumptions will be subjective, and our answers will depend on them. Exactly the same can be said for the other probabilities in our generative model too. We are assuming, for example, that the balls are drawn from an urn independently; but could there not be correlations in the sequence because Fred’s ball-drawing action is not perfectly random? Indeed there could be, so the likelihood function that we use depends on assumptions too. In real data modelling problems, priors are subjective and so are likelihoods. We are now using P () to denote probability densities over continuous variables as well as probabilities over discrete variables and probabilities of logical propositions. The probability that a continuous variable v lies between values Rb a and b (where b > a) is defined to be a dv P (v). P (v)dv is dimensionless. The density P (v) is a dimensional quantity, having dimensions inverse to the dimensions of v – in contrast to discrete probabilities, which are dimensionless. Don’t be surprised to see probability densities greater than 1. This is normal, Rb and nothing is wrong, as long as a dv P (v) ≤ 1 for any interval (a, b). Conditional and joint probability densities are defined in just the same way as conditional and joint probabilities.

. Exercise 2.8.[2 ] Assuming a uniform prior on fH , P (fH ) = 1, solve the problem posed in example 2.7 (p.30). Sketch the posterior distribution of f H and compute the probability that the N +1th outcome will be a head, for (a) N = 3 and nH = 0; (b) N = 3 and nH = 2; (c) N = 10 and nH = 3; (d) N = 300 and nH = 29. You will find the beta integral useful: Z 1 Fa !Fb ! Γ(Fa + 1)Γ(Fb + 1) = . (2.32) dpa pFa a (1 − pa )Fb = Γ(F + F + 2) (F + Fb + 1)! a a b 0

Here P (f ) denotes a probability density, rather than a probability distribution.

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2.3: Forward probabilities and inverse probabilities

31

You may also find it instructive to look back at example 2.6 (p.27) and equation (2.31). People sometimes confuse assigning a prior distribution to an unknown parameter such as fH with making an initial guess of the value of the parameter. But the prior over fH , P (fH ), is not a simple statement like ‘initially, I would guess fH = 1/2’. The prior is a probability density over f H which specifies the prior degree of belief that fH lies in any interval (f, f + δf ). It may well be the case that our prior for fH is symmetric about 1/2, so that the mean of fH under the prior is 1/2. In this case, the predictive distribution for the first toss x1 would indeed be Z Z P (x1 = head) = dfH P (fH )P (x1 = head | fH ) = dfH P (fH )fH = 1/2.

(2.33) But the prediction for subsequent tosses will depend on the whole prior distribution, not just its mean. Data compression and inverse probability Consider the following task. Example 2.9. Write a computer program capable of compressing binary files like this one:

0000000000000000000010010001000000100000010000000000000000000000000000000000001010000000000000110000 1000000000010000100000000010000000000000000000000100000000000000000100000000011000001000000011000100 0000000001001000000000010001000000000000000011000000000000000000000000000010000000000000000100000000

The string shown contains n1 = 29 1s and n0 = 271 0s. Intuitively, compression works by taking advantage of the predictability of a file. In this case, the source of the file appears more likely to emit 0s than 1s. A data compression program that compresses this file must, implicitly or explicitly, be addressing the question ‘What is the probability that the next character in this file is a 1?’ Do you think this problem is similar in character to example 2.7 (p.30)? I do. One of the themes of this book is that data compression and data modelling are one and the same, and that they should both be addressed, like the urn of example 2.6, using inverse probability. Example 2.9 is solved in Chapter 6.

The likelihood principle Please solve the following two exercises. Example 2.10. Urn A contains three balls: one black, and two white; urn B contains three balls: two black, and one white. One of the urns is selected at random and one ball is drawn. The ball is black. What is the probability that the selected urn is urn A? Example 2.11. Urn A contains five balls: one black, two white, one green and one pink; urn B contains five hundred balls: two hundred black, one hundred white, 50 yellow, 40 cyan, 30 sienna, 25 green, 25 silver, 20 gold, and 10 purple. [One fifth of A’s balls are black; two-fifths of B’s are black.] One of the urns is selected at random and one ball is drawn. The ball is black. What is the probability that the urn is urn A?

A

B

Figure 2.7. Urns for example 2.10.

c g p

y

... ... ...

g p s

Figure 2.8. Urns for example 2.11.

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32

2 — Probability, Entropy, and Inference

What do you notice about your solutions? Does each answer depend on the detailed contents of each urn? The details of the other possible outcomes and their probabilities are irrelevant. All that matters is the probability of the outcome that actually happened (here, that the ball drawn was black) given the different hypotheses. We need only to know the likelihood, i.e., how the probability of the data that happened varies with the hypothesis. This simple rule about inference is known as the likelihood principle. The likelihood principle: given a generative model for data d given parameters θ, P (d | θ), and having observed a particular outcome d1 , all inferences and predictions should depend only on the function P (d1 | θ). In spite of the simplicity of this principle, many classical statistical methods violate it.

2.4 Definition of entropy and related functions The Shannon information content of an outcome x is defined to be h(x) = log 2

1 . P (x)

(2.34)

It is measured in bits. [The word ‘bit’ is also used to denote a variable whose value is 0 or 1; I hope context will always make clear which of the two meanings is intended.] In the next few chapters, we will establish that the Shannon information content h(ai ) is indeed a natural measure of the information content of the event x = ai . At that point, we will shorten the name of this quantity to ‘the information content’. The fourth column in table 2.9 shows the Shannon information content of the 27 possible outcomes when a random character is picked from an English document. The outcome x = z has a Shannon information content of 10.4 bits, and x = e has an information content of 3.5 bits. The entropy of an ensemble X is defined to be the average Shannon information content of an outcome: X 1 H(X) ≡ P (x) log , (2.35) P (x) x∈AX

with the convention limθ→0+ θ log 1/θ = 0.

for

P (x) = 0

that

0 × log 1/0 ≡ 0,

since

Like the information content, entropy is measured in bits. When it is convenient, we may also write H(X) as H(p), where p is the vector (p1 , p2 , . . . , pI ). Another name for the entropy of X is the uncertainty of X. Example 2.12. The entropy of a randomly selected letter in an English document is about 4.11 bits, assuming its probability is as given in table 2.9. We obtain this number by averaging log 1/p i (shown in the fourth column) under the probability distribution p i (shown in the third column).

i

ai

pi

h(pi )

1 2 3 4 5 6 7 8 9 10 11 12 13 14 15 16 17 18 19 20 21 22 23 24 25 26 27

a b c d e f g h i j k l m n o p q r s t u v w x y z -

.0575 .0128 .0263 .0285 .0913 .0173 .0133 .0313 .0599 .0006 .0084 .0335 .0235 .0596 .0689 .0192 .0008 .0508 .0567 .0706 .0334 .0069 .0119 .0073 .0164 .0007 .1928

4.1 6.3 5.2 5.1 3.5 5.9 6.2 5.0 4.1 10.7 6.9 4.9 5.4 4.1 3.9 5.7 10.3 4.3 4.1 3.8 4.9 7.2 6.4 7.1 5.9 10.4 2.4

1 pi

4.1

X i

pi log2

Table 2.9. Shannon information contents of the outcomes a–z.

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2.5: Decomposability of the entropy

33

We now note some properties of the entropy function. • H(X) ≥ 0 with equality iff pi = 1 for one i. [‘iff’ means ‘if and only if’.] • Entropy is maximized if p is uniform: H(X) ≤ log(|AX |) with equality iff pi = 1/|AX | for all i.

(2.36)

Notation: the vertical bars ‘| · |’ have two meanings. If A X is a set, |AX | denotes the number of elements in AX ; if x is a number, then |x| is the absolute value of x. The redundancy measures the fractional difference between H(X) and its maximum possible value, log(|AX |). The redundancy of X is: 1−

H(X) . log |AX |

(2.37)

We won’t make use of ‘redundancy’ in this book, so I have not assigned a symbol to it. The joint entropy of X, Y is: H(X, Y ) =

X

xy∈AX AY

P (x, y) log

1 . P (x, y)

(2.38)

Entropy is additive for independent random variables: H(X, Y ) = H(X) + H(Y ) iff P (x, y) = P (x)P (y).

(2.39)

Our definitions for information content so far apply only to discrete probability distributions over finite sets AX . The definitions can be extended to infinite sets, though the entropy may then be infinite. The case of a probability density over a continuous set is addressed in section 11.3. Further important definitions and exercises to do with entropy will come along in section 8.1.

2.5 Decomposability of the entropy The entropy function satisfies a recursive property that can be very useful when computing entropies. For convenience, we’ll stretch our notation so that we can write H(X) as H(p), where p is the probability vector associated with the ensemble X. Let’s illustrate the property by an example first. Imagine that a random variable x ∈ {0, 1, 2} is created by first flipping a fair coin to determine whether x = 0; then, if x is not 0, flipping a fair coin a second time to determine whether x is 1 or 2. The probability distribution of x is 1 1 1 P (x = 0) = ; P (x = 1) = ; P (x = 2) = . 2 4 4

(2.40)

What is the entropy of X? We can either compute it by brute force: H(X) = 1/2 log 2 + 1/4 log 4 + 1/4 log 4 = 1.5;

(2.41)

or we can use the following decomposition, in which the value of x is revealed gradually. Imagine first learning whether x = 0, and then, if x is not 0, learning which non-zero value is the case. The revelation of whether x = 0 or not entails

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34

2 — Probability, Entropy, and Inference

revealing a binary variable whose probability distribution is { 1/2, 1/2}. This revelation has an entropy H(1/2, 1/2) = 12 log 2 + 12 log 2 = 1 bit. If x is not 0, we learn the value of the second coin flip. This too is a binary variable whose probability distribution is {1/2, 1/2}, and whose entropy is 1 bit. We only get to experience the second revelation half the time, however, so the entropy can be written: H(X) = H(1/2, 1/2) + 1/2 H(1/2, 1/2). (2.42) Generalizing, the observation we are making about the entropy of any probability distribution p = {p1 , p2 , . . . , pI } is that   p3 pI p2 , ,..., . (2.43) H(p) = H(p1 , 1−p1 ) + (1−p1 )H 1−p1 1−p1 1−p1 When it’s written as a formula, this property looks regrettably ugly; nevertheless it is a simple property and one that you should make use of. Generalizing further, the entropy has the property for any m that H(p) = H [(p1 + p2 + · · · + pm ), (pm+1 + pm+2 + · · · + pI )]   pm p1 ,..., +(p1 + · · · + pm )H (p1 + · · · + pm ) (p1 + · · · + pm )   pI pm+1 +(pm+1 + · · · + pI )H ,..., . (pm+1 + · · · + pI ) (pm+1 + · · · + pI ) (2.44) Example 2.13. A source produces a character x from the alphabet A = {0, 1, . . . , 9, a, b, . . . , z}; with probability 1/3, x is a numeral (0, . . . , 9); with probability 1/3, x is a vowel (a, e, i, o, u); and with probability 1/3 it’s one of the 21 consonants. All numerals are equiprobable, and the same goes for vowels and consonants. Estimate the entropy of X. Solution. log 3 + 13 (log 10 + log 5 + log 21) = log 3 +

1 3

log 1050 ' log 30 bits. 2

2.6 Gibbs’ inequality The relative entropy or Kullback–Leibler divergence between two probability distributions P (x) and Q(x) that are defined over the same alphabet AX is DKL (P ||Q) =

X

P (x) log

x

P (x) . Q(x)

(2.45)

The relative entropy satisfies Gibbs’ inequality DKL (P ||Q) ≥ 0

(2.46)

with equality only if P = Q. Note that in general the relative entropy is not symmetric under interchange of the distributions P and Q: in general DKL (P ||Q) 6= DKL (Q||P ), so DKL , although it is sometimes called the ‘KL distance’, is not strictly a distance. The relative entropy is important in pattern recognition and neural networks, as well as in information theory. Gibbs’ inequality is probably the most important inequality in this book. It, and many other inequalities, can be proved using the concept of convexity.

The ‘ei’ in Leibler is pronounced the same as in heist.

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2.7: Jensen’s inequality for convex functions

35

2.7 Jensen’s inequality for convex functions The words ‘convex ^’ and ‘concave _’ may be pronounced ‘convex-smile’ and ‘concave-frown’. This terminology has useful redundancy: while one may forget which way up ‘convex’ and ‘concave’ are, it is harder to confuse a smile with a frown.

Convex ^ functions. A function f (x) is convex ^ over (a, b) if every chord of the function lies above the function, as shown in figure 2.10; that is, for all x1 , x2 ∈ (a, b) and 0 ≤ λ ≤ 1, f (λx1 + (1 − λ)x2 ) ≤

λf (x1 ) + (1 − λ)f (x2 ).

(2.47)

A function f is strictly convex ^ if, for all x 1 , x2 ∈ (a, b), the equality holds only for λ = 0 and λ = 1.

λf (x1 ) + (1 − λ)f (x2 ) f (x∗ ) x1

x2 x∗ = λx1 + (1 − λ)x2

Figure 2.10. Definition of convexity.

Similar definitions apply to concave _ and strictly concave _ functions. Some strictly convex ^ functions are • x2 , ex and e−x for all x; • log(1/x) and x log x for x > 0.

-1

0

1

log x1

e−x

x2

2

3

-1

0

1

2

3

0

1

x log x

2

3 0

1

2

Figure 2.11. Convex ^ functions.

3

Jensen’s inequality. If f is a convex ^ function and x is a random variable then: E [f (x)] ≥ f (E[x]) , (2.48) where E denotes expectation. If f is strictly convex ^ and E [f (x)] = f (E[x]), then the random variable x is a constant. Jensen’s inequality can also be rewritten for a concave _ function, with the direction of the inequality reversed. A physical version of Jensen’s inequality runs as follows. If a collection of masses pi are placed on a convex ^ curve f (x) at locations (xi , f (xi )), then the centre of gravity of those masses, which is at (E[x], E [f (x)]), lies above the curve. If this fails to convince you, then feel free to do the following exercise. Exercise 2.14.[2, p.41] Prove Jensen’s inequality. Example 2.15. Three squares have average area A¯ = 100 m2 . The average of the lengths of their sides is ¯l = 10 m. What can be said about the size of the largest of the three squares? [Use Jensen’s inequality.] Solution. Let x be the length of the side of a square, and let the probability of x be 1/3, 1/3, 1/3 over the three lengths l1 , l2 , l3 . Then the information that we have is that E [x] = 10 and E [f (x)] = 100, where f (x) = x 2 is the function mapping lengths to areas. This is a strictly convex ^ function. We notice that the equality E [f (x)] = f (E[x]) holds, therefore x is a constant, and the three lengths must all be equal. The area of the largest square is 100 m 2 . 2

Centre of gravity

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36

2 — Probability, Entropy, and Inference

Convexity and concavity also relate to maximization If f (x) is concave _ and there exists a point at which ∂f = 0 for all k, ∂xk

(2.49)

then f (x) has its maximum value at that point. The converse does not hold: if a concave _ f (x) is maximized at some x it is not necessarily true that the gradient ∇f (x) is equal to zero there. For example, f (x) = −|x| is maximized at x = 0 where its derivative is undefined; and f (p) = log(p), for a probability p ∈ (0, 1), is maximized on the boundary of the range, at p = 1, where the gradient df (p)/dp = 1.

2.8 Exercises Sums of random variables Exercise 2.16.[3, p.41] (a) Two ordinary dice with faces labelled 1, . . . , 6 are thrown. What is the probability distribution of the sum of the values? What is the probability distribution of the absolute difference between the values? (b) One hundred ordinary dice are thrown. What, roughly, is the probability distribution of the sum of the values? Sketch the probability distribution and estimate its mean and standard deviation. (c) How can two cubical dice be labelled using the numbers {0, 1, 2, 3, 4, 5, 6} so that when the two dice are thrown the sum has a uniform probability distribution over the integers 1–12? (d) Is there any way that one hundred dice could be labelled with integers such that the probability distribution of the sum is uniform?

Inference problems Exercise 2.17.[2, p.41] If q = 1 − p and a = ln p/q, show that p=

1 . 1 + exp(−a)

(2.50)

Sketch this function and find its relationship to the hyperbolic tangent u −u function tanh(u) = eeu −e +e−u . It will be useful to be fluent in base-2 logarithms also. If b = log 2 p/q, what is p as a function of b? . Exercise 2.18.[2, p.42] Let x and y be dependent random variables with x a binary variable taking values in AX = {0, 1}. Use Bayes’ theorem to show that the log posterior probability ratio for x given y is log

P (x = 1 | y) P (y | x = 1) P (x = 1) = log + log . P (x = 0 | y) P (y | x = 0) P (x = 0)

(2.51)

. Exercise 2.19.[2, p.42] Let x, d1 and d2 be random variables such that d1 and d2 are conditionally independent given a binary variable x. Use Bayes’ theorem to show that the posterior probability ratio for x given {d i } is P (d1 | x = 1) P (d2 | x = 1) P (x = 1) P (x = 1 | {di }) = . P (x = 0 | {di }) P (d1 | x = 0) P (d2 | x = 0) P (x = 0)

(2.52)

This exercise is intended to help you think about the central-limit theorem, which says that if independent random variables x1 , x2 , . . . , xN have means µn and finite variances σn2 , then, in Pthe limit of large N , the sum n xn has a distribution that tends to a normal (Gaussian) distribution P with n µn and variance P mean 2 n σn .

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2.8: Exercises

37

Life in high-dimensional spaces Probability distributions and volumes have some unexpected properties in high-dimensional spaces. Exercise 2.20.[2, p.42] Consider a sphere of radius r in an N -dimensional real space. Show that the fraction of the volume of the sphere that is in the surface shell lying at values of the radius between r −  and r, where 0 <  < r, is:   N f =1− 1− . (2.53) r Evaluate f for the cases N = 2, N = 10 and N = 1000, with (a) /r = 0.01; (b) /r = 0.5. Implication: points that are uniformly distributed in a sphere in N dimensions, where N is large, are very likely to be in a thin shell near the surface.

Expectations and entropies You are probably familiar with the idea of computing the expectation of a function of x, X E [f (x)] = hf (x)i = P (x)f (x). (2.54) x

Maybe you are not so comfortable with computing this expectation in cases where the function f (x) depends on the probability P (x). The next few examples address this concern.

Exercise 2.21.[1, p.43] Let pa = 0.1, pb = 0.2, and pc = 0.7. Let f (a) = 10, f (b) = 5, and f (c) = 10/7. What is E [f (x)]? What is E [1/P (x)]? Exercise 2.22.[2, p.43] For an arbitrary ensemble, what is E [1/P (x)]? . Exercise 2.23.[1, p.43] Let pa = 0.1, pb = 0.2, and pc = 0.7. Let g(a) = 0, g(b) = 1, and g(c) = 0. What is E [g(x)]? . Exercise 2.24.[1, p.43] Let pa = 0.1, pb = 0.2, and pc = 0.7. What is the probability that P (x) ∈ [0.15, 0.5]? What is   P (x) P log > 0.05 ? 0.2 Exercise 2.25.[3, p.43] Prove the assertion that H(X) ≤ log(|A X |) with equality iff pi = 1/|AX | for all i. (|AX | denotes the number of elements in the set AX .) [Hint: use Jensen’s inequality (2.48); if your first attempt to use Jensen does not succeed, remember that Jensen involves both a random variable and a function, and you have quite a lot of freedom in choosing these; think about whether your chosen function f should be convex or concave.]

. Exercise 2.26.[3, p.44] Prove that the relative entropy (equation (2.45)) satisfies DKL (P ||Q) ≥ 0 (Gibbs’ inequality) with equality only if P = Q. . Exercise 2.27.[2 ] Prove that the entropy is indeed decomposable as described in equations (2.43–2.44).

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38

2 — Probability, Entropy, and Inference

. Exercise 2.28.[2, p.45] A random variable x ∈ {0, 1, 2, 3} is selected by flipping a bent coin with bias f to determine whether the outcome is in {0, 1} or {2, 3}; then either flipping a second bent coin with bias g or a third bent coin with bias h respectively. Write down the probability distribution of x. Use the decomposability of the entropy (2.44) to find the entropy of X. [Notice how compact an expression is obtained if you make use of the binary entropy function H2 (x), compared with writing out the four-term entropy explicitly.] Find the derivative of H(X) with respect to f . [Hint: dH2 (x)/dx = log((1 − x)/x).] . Exercise 2.29.[2, p.45] An unbiased coin is flipped until one head is thrown. What is the entropy of the random variable x ∈ {1, 2, 3, . . .}, the number of flips? Repeat the calculation for the case of a biased coin with probability f of coming up heads. [Hint: solve the problem both directly and by using the decomposability of the entropy (2.43).]

2.9 Further exercises Forward probability . Exercise 2.30.[1 ] An urn contains w white balls and b black balls. Two balls are drawn, one after the other, without replacement. Prove that the probability that the first ball is white is equal to the probability that the second is white. . Exercise 2.31.[2 ] A circular coin of diameter a is thrown onto a square grid whose squares are b × b. (a < b) What is the probability that the coin will lie entirely within one square? [Ans: (1 − a/b) 2 ] . Exercise 2.32.[3 ] Buffon’s needle. A needle of length a is thrown onto a plane covered with equally spaced parallel lines with separation b. What is the probability that the needle will cross a line? [Ans, if a < b: 2a/πb] [Generalization – Buffon’s noodle: on average, a random curve of length A is expected to intersect the lines 2A/πb times.] Exercise 2.33.[2 ] Two points are selected at random on a straight line segment of length 1. What is the probability that a triangle can be constructed out of the three resulting segments? Exercise 2.34.[2, p.45] An unbiased coin is flipped until one head is thrown. What is the expected number of tails and the expected number of heads? Fred, who doesn’t know that the coin is unbiased, estimates the bias using fˆ ≡ h/(h + t), where h and t are the numbers of heads and tails ˆ tossed. Compute and sketch the probability distribution of f. N.B., this is a forward probability problem, a sampling theory problem, not an inference problem. Don’t use Bayes’ theorem. Exercise 2.35.[2, p.45] Fred rolls an unbiased six-sided die once per second, noting the occasions when the outcome is a six. (a) What is the mean number of rolls from one six to the next six? (b) Between two rolls, the clock strikes one. What is the mean number of rolls until the next six?

g 0  f

@ 1−gR @ 1

@ 1−f @ h 2 R @  @ @ 3 1−hR

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2.9: Further exercises

39

(c) Now think back before the clock struck. What is the mean number of rolls, going back in time, until the most recent six? (d) What is the mean number of rolls from the six before the clock struck to the next six? (e) Is your answer to (d) different from your answer to (a)? Explain. Another version of this exercise refers to Fred waiting for a bus at a bus-stop in Poissonville where buses arrive independently at random (a Poisson process), with, on average, one bus every six minutes. What is the average wait for a bus, after Fred arrives at the stop? [6 minutes.] So what is the time between the two buses, the one that Fred just missed, and the one that he catches? [12 minutes.] Explain the apparent paradox. Note the contrast with the situation in Clockville, where the buses are spaced exactly 6 minutes apart. There, as you can confirm, the mean wait at a bus-stop is 3 minutes, and the time between the missed bus and the next one is 6 minutes.

Conditional probability . Exercise 2.36.[2 ] You meet Fred. Fred tells you he has two brothers, Alf and Bob. What is the probability that Fred is older than Bob? Fred tells you that he is older than Alf. Now, what is the probability that Fred is older than Bob? (That is, what is the conditional probability that F > B given that F > A?) . Exercise 2.37.[2 ] The inhabitants of an island tell the truth one third of the time. They lie with probability 2/3. On an occasion, after one of them made a statement, you ask another ‘was that statement true?’ and he says ‘yes’. What is the probability that the statement was indeed true? . Exercise 2.38.[2, p.46] Compare two ways of computing the probability of error of the repetition code R3 , assuming a binary symmetric channel (you did this once for exercise 1.2 (p.7)) and confirm that they give the same answer. Binomial distribution method. Add the probability that all three bits are flipped to the probability that exactly two bits are flipped. Sum rule method. Using the sum rule, compute the marginal probability that P r takes on each of the eight possible values, P (r). [P (r) = s P (s)P (r | s).] Then compute the posterior probability of s for each of the eight values of r. [In fact, by symmetry, only two example cases r = (000) and r = (001) need be considered.] Notice that some of the inferred bits are better determined than others. From the posterior probability P (s | r) you can read out the case-by-case error probability, the probability that the more probable hypothesis is not correct, P (error | r). Find the average error probability using the sum rule, X P (error) = P (r)P (error | r). (2.55) r

Equation (1.18) gives the posterior probability of the input s, given the received vector r.

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40

2 — Probability, Entropy, and Inference

. Exercise 2.39.[3C, p.46] The frequency pn of the nth most frequent word in English is roughly approximated by 

pn '

0.1 n

0

for n ∈ 1, . . . , 12 367 n > 12 367.

(2.56)

[This remarkable 1/n law is known as Zipf’s law, and applies to the word frequencies of many languages (Zipf, 1949).] If we assume that English is generated by picking words at random according to this distribution, what is the entropy of English (per word)? [This calculation can be found in ‘Prediction and entropy of printed English’, C.E. Shannon, Bell Syst. Tech. J. 30, pp.50–64 (1950), but, inexplicably, the great man made numerical errors in it.]

2.10 Solutions Solution to exercise 2.2 (p.24). No, they are not independent. If they were then all the conditional distributions P (y | x) would be identical functions of y, regardless of x (cf. figure 2.3). Solution to exercise 2.4 (p.27).

We define the fraction f B ≡ B/K.

(a) The number of black balls has a binomial distribution. P (nB | fB , N ) =



 N f nB (1 − fB )N −nB . nB B

(2.57)

(b) The mean and variance of this distribution are: E[nB ] = N fB

(2.58)

var[nB ] = N fB (1 − fB ).

(2.59)

These results p in example 1.1 (p.1). The standard deviation p were derived of nB is var[nB ] = N fB (1 − fB ).

When B/K = 1/5 and N = 5, the expectation and variance of n B are 1 and 4/5. The standard deviation is 0.89. When B/K = 1/5 and N = 400, the expectation and variance of n B are 80 and 64. The standard deviation is 8. Solution to exercise 2.5 (p.27).

The numerator of the quantity

z=

(nB − fB N )2 N fB (1 − fB )

can be recognized as (nB − E[nB ])2 ; the denominator is equal to the variance of nB (2.59), which is by definition the expectation of the numerator. So the expectation of z is 1. [A random variable like z, which measures the deviation of data from the expected value, is sometimes called χ 2 (chi-squared).] In the case N = 5 and fB = 1/5, N fB is 1, and var[nB ] is 4/5. The numerator has five possible values, only one of which is smaller than 1: (n B − fB N )2 = 0 has probability P (nB = 1) = 0.4096; so the probability that z < 1 is 0.4096.

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2.10: Solutions

41

Solution to exercise 2.14 (p.35).

We wish to prove, given the property

f (λx1 + (1 − λ)x2 ) ≤ λf (x1 ) + (1 − λ)f (x2 ), that, if

P

(2.60)

pi = 1 and pi ≥ 0, I X i=1

I X

pi f (xi ) ≥ f

pi xi

i=1

!

.

(2.61)

We proceed by recursion, working from the right-hand side. (This proof does not handle cases where some pi = 0; such details are left to the pedantic reader.) At the first line we use the definition of convexity (2.60) with λ = p1 = p1 ; at the second line, λ = Ip2 . I i=1

f

pi

I X i=1

i=2

pi xi

!

=f

≤ p1 f (x1 ) + ≤ p1 f (x1 ) +

p1 x1 +

I X

pi xi

i=2

"

I X

i=2 " I X i=2

pi pi

#" #"

and so forth.

I X

f

,

pi xi

i=2

p2 PI

i=2 pi

pi

! I X

pi

i=2

f (x2 ) +

!#

PI

pi f Pi=3 I i=2 pi

(2.62) I X i=3

pi xi

,

I X i=3

pi

!#

, 2

Solution to exercise 2.16 (p.36). (a) For the outcomes {2, 3, 4, 5, 6, 7, 8, 9, 10, 11, 12}, the probabilities are P = 1 2 3 4 5 6 5 4 3 2 1 , 36 , 36 , 36 , 36 , 36 , 36 , 36 , 36 , 36 , 36 }. { 36 (b) The value of one die has mean 3.5 and variance 35/12. So the sum of one hundred has mean 350 and variance 3500/12 ' 292, and by the central-limit theorem the probability distribution is roughly Gaussian (but confined to the integers), with this mean and variance. (c) In order to obtain a sum that has a uniform distribution we have to start from random variables some of which have a spiky distribution with the probability mass concentrated at the extremes. The unique solution is to have one ordinary die and one with faces 6, 6, 6, 0, 0, 0. (d) Yes, a uniform distribution can be created in several ways, for example by labelling the rth die with the numbers {0, 1, 2, 3, 4, 5} × 6 r . Solution to exercise 2.17 (p.36).

and q = 1 − p gives

a = ln

p q

p 1−p

= ea



p =

p = ea q



(2.63)

(2.64) ea

ea

+1

=

1 . 1 + exp(−a)

(2.65)

The hyperbolic tangent is tanh(a) =

ea − e−a ea + e−a

(2.66)

To think about: does this uniform distribution contradict the central-limit theorem?

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42

2 — Probability, Entropy, and Inference

so f (a) ≡ =

 1 − e−a + 1 1 + e−a ! 1 + 1 = (tanh(a/2) + 1). 2

1 1 = 1 + exp(−a) 2 ea/2 − e−a/2 ea/2 + e−a/2

1 2



(2.67)

In the case b = log 2 p/q, we can repeat steps (2.63–2.65), replacing e by 2, to obtain 1 . (2.68) p= 1 + 2−b Solution to exercise 2.18 (p.36). P (y | x)P (x) P (y) P (y | x = 1) P (x = 1) = P (y | x = 0) P (x = 0) P (y | x = 1) P (x = 1) = log + log . P (y | x = 0) P (x = 0)

P (x | y) = P (x = 1 | y) P (x = 0 | y) P (x = 1 | y) ⇒ log P (x = 0 | y) ⇒

(2.69) (2.70) (2.71)

Solution to exercise 2.19 (p.36). The conditional independence of d 1 and d2 given x means P (x, d1 , d2 ) = P (x)P (d1 | x)P (d2 | x). (2.72) This gives a separation of the posterior probability ratio into a series of factors, one for each data point, times the prior probability ratio. P (x = 1 | {di }) P (x = 0 | {di })

= =

P ({di } | x = 1) P (x = 1) P ({di } | x = 0) P (x = 0) P (d1 | x = 1) P (d2 | x = 1) P (x = 1) . P (d1 | x = 0) P (d2 | x = 0) P (x = 0)

(2.73) (2.74)

Life in high-dimensional spaces Solution to exercise 2.20 (p.37). N dimensions is in fact

The volume of a hypersphere of radius r in

V (r, N ) =

π N/2 N r , (N/2)!

(2.75)

but you don’t need to know this. For this question all that we need is the r-dependence, V (r, N ) ∝ r N . So the fractional volume in (r − , r) is   N r N − (r − )N = 1 − 1 − . rN r

(2.76)

The fractional volumes in the shells for the required cases are: N

2

10

1000

/r = 0.01 /r = 0.5

0.02 0.75

0.096 0.999

0.99996 1 − 2−1000

Notice that no matter how small  is, for large enough N essentially all the probability mass is in the surface shell of thickness .

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2.10: Solutions

43

Solution to exercise 2.21 (p.37). f (b) = 5, and f (c) = 10/7.

p a = 0.1, pb = 0.2, pc = 0.7.

E [f (x)] = 0.1 × 10 + 0.2 × 5 + 0.7 × 10/7 = 3.

f (a) = 10, (2.77)

For each x, f (x) = 1/P (x), so E [1/P (x)] = E [f (x)] = 3.

(2.78)

Solution to exercise 2.22 (p.37). For general X, X X E [1/P (x)] = P (x)1/P (x) = 1 = |AX |. x∈AX

(2.79)

x∈AX

Solution to exercise 2.23 (p.37). p a = 0.1, pb = 0.2, pc = 0.7. g(a) = 0, g(b) = 1, and g(c) = 0. E [g(x)] = pb = 0.2. (2.80) Solution to exercise 2.24 (p.37). P (P (x) ∈ [0.15, 0.5]) = pb = 0.2.   P (x) > 0.05 = pa + pc = 0.8. P log 0.2

(2.81) (2.82)

Solution to exercise 2.25 (p.37). This type of question can be approached in two ways: either by differentiating the function to be maximized, finding the maximum, and proving it is a global maximum; this strategy is somewhat risky since it is possible for the maximum of a function to be at the boundary of the space, at a place where the derivative is not zero. Alternatively, a carefully chosen inequality can establish the answer. The second method is much neater. Proof by differentiation (not the recommended method). Since it is slightly easier to differentiate ln 1/p than log 2 1/p, we temporarily define H(X) to be measured using natural logarithms, thus scaling it down by a factor of log 2 e. X 1 H(X) = pi ln (2.83) pi i

∂H(X) ∂pi

= ln

we maximize subject to the constraint a Lagrange multiplier:

P

1 −1 pi

i pi

G(p) ≡ H(X) + λ ∂G(p) ∂pi

= ln

(2.84)

= 1 which can be enforced with X i

pi − 1

1 − 1 + λ. pi

!

(2.85) (2.86)

At a maximum, ln

1 −1+λ = 0 pi 1 = 1 − λ, ⇒ ln pi

(2.87) (2.88)

so all the pi are equal. That this extremum is indeed a maximum is established by finding the curvature: 1 ∂ 2 G(p) = − δij , (2.89) ∂pi ∂pj pi which is negative definite. 2

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44

2 — Probability, Entropy, and Inference

Proof using Jensen’s inequality (recommended method). the inequality.

First a reminder of

If f is a convex ^ function and x is a random variable then: E [f (x)] ≥ f (E[x]) . If f is strictly convex ^ and E [f (x)] = f (E[x]), then the random variable x is a constant (with probability 1). The secret of a proof using Jensen’s inequality is to choose the right function and the right random variable. We could define f (u) = log

1 = − log u u

(2.90)

P (which is a convex function) and think of H(X) = pi log p1i as the mean of f (u) where u = P (x), but this would not get us there – it would give us an inequality in the wrong direction. If instead we define u = 1/P (x)

(2.91)

H(X) = −E [f (1/P (x))] ≤ −f (E[1/P (x)]) ;

(2.92)

then we find:

now we know from exercise 2.22 (p.37) that E[1/P (x)] = |A X |, so H(X) ≤ −f (|AX |) = log |AX |.

(2.93)

Equality holds only if the random variable u = 1/P (x) is a constant, which means P (x) is a constant for all x. 2 Solution to exercise 2.26 (p.37). DKL (P ||Q) =

X x

P (x) log

P (x) . Q(x)

(2.94)

We prove Gibbs’ inequality using Jensen’s inequality. Let f (u) = log 1/u and . Then u = Q(x) P (x) DKL (P ||Q) = E[f (Q(x)/P (x))] !   X Q(x) 1 ≥ f P (x) = log P = 0, P (x) x Q(x) x

with equality only if u =

Q(x) P (x)

is a constant, that is, if Q(x) = P (x).

(2.95) (2.96) 2

Second solution. In the above proof the expectations were with respect to the probability distribution P (x). A second solution method uses Jensen’s (x) inequality with Q(x) instead. We define f (u) = u log u and let u = PQ(x) . Then   X P (x) P (x) P (x) X Q(x) DKL (P ||Q) = Q(x)f log = (2.97) Q(x) Q(x) Q(x) x x ! X P (x) ≥ f Q(x) = f (1) = 0, (2.98) Q(x) x with equality only if u =

P (x) Q(x)

is a constant, that is, if Q(x) = P (x).

2

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2.10: Solutions

45

Solution to exercise 2.28 (p.38). H(X) = H2 (f ) + f H2 (g) + (1 − f )H2 (h).

(2.99)

Solution to exercise 2.29 (p.38). The probability that there are x − 1 tails and then one head (so we get the first head on the xth toss) is P (x) = (1 − f )x−1 f.

(2.100)

If the first toss is a tail, the probability distribution for the future looks just like it did before we made the first toss. Thus we have a recursive expression for the entropy: H(X) = H2 (f ) + (1 − f )H(X). (2.101) Rearranging, H(X) = H2 (f )/f. Solution to exercise 2.34 (p.38).

(2.102)

The probability of the number of tails t is

 t 1 1 for t ≥ 0. P (t) = 2 2

(2.103)

The expected number of heads is 1, by definition of the problem. The expected number of tails is ∞  t X 1 1 t E[t] = , (2.104) 2 2 t=0

which may be shown to be 1 in a variety of ways. For example, since the situation after one tail is thrown is equivalent to the opening situation, we can write down the recurrence relation 1 1 E[t] = (1 + E[t]) + 0 ⇒ E[t] = 1. 2 2

ˆ P (f)

0.4 0.3 0.2 0.1 0 0

(2.105)

The probability distribution of the ‘estimator’ fˆ = 1/(1 + t), given that f = 1/2, is plotted in figure 2.12. The probability of fˆ is simply the probability of the corresponding value of t. Solution to exercise 2.35 (p.38). (a) The mean number of rolls from one six to the next six is six (assuming we start counting rolls after the first of the two sixes). The probability that the next six occurs on the rth roll is the probability of not getting a six for r − 1 rolls multiplied by the probability of then getting a six:  r−1 1 5 , for r ∈ {1, 2, 3, . . .}. (2.106) P (r1 = r) = 6 6 This probability distribution of the number of rolls, r, may be called an exponential distribution, since P (r1 = r) = e−αr /Z,

0.5

(2.107)

where α = ln(6/5), and Z is a normalizing constant. (b) The mean number of rolls from the clock until the next six is six. (c) The mean number of rolls, going back in time, until the most recent six is six.

0.2

0.4

0.6

0.8

1

fˆ Figure 2.12. The probability distribution of the estimator fˆ = 1/(1 + t), given that f = 1/2.

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46

2 — Probability, Entropy, and Inference

(d) The mean number of rolls from the six before the clock struck to the six after the clock struck is the sum of the answers to (b) and (c), less one, that is, eleven. (e) Rather than explaining the difference between (a) and (d), let me give another hint. Imagine that the buses in Poissonville arrive independently at random (a Poisson process), with, on average, one bus every six minutes. Imagine that passengers turn up at bus-stops at a uniform rate, and are scooped up by the bus without delay, so the interval between two buses remains constant. Buses that follow gaps bigger than six minutes become overcrowded. The passengers’ representative complains that two-thirds of all passengers found themselves on overcrowded buses. The bus operator claims, ‘no, no – only one third of our buses are overcrowded’. Can both these claims be true? Solution to exercise 2.38 (p.39). Binomial distribution method. From the solution to exercise 1.2, p B = 3f 2 (1 − f ) + f 3 . Sum rule method. The marginal probabilities of the eight values of r are illustrated by: P (r = 000) = 1/2(1 − f )3 + 1/2f 3 , (2.108) P (r = 001) = 1/2f (1 − f )2 + 1/2f 2 (1 − f ) = 1/2f (1 − f ).

(2.109)

The posterior probabilities are represented by P (s = 1 | r = 000) = and P (s = 1 | r = 001) =

f3 (1 − f )3 + f 3

(1 − f )f 2 = f. f (1 − f )2 + f 2 (1 − f )

0.15 0.1 0.05 0 0

5

10

15

20

Figure 2.13. The probability distribution of the number of rolls r1 from one 6 to the next (falling solid line), P (r1 = r) =

 r−1 1 5 , 6 6

and the probability distribution (dashed line) of the number of rolls from the 6 before 1pm to the next 6, rtot , P (rtot = r) = r

 r−1  2 5 1 . 6 6

The probability P (r1 > 6) is about 1/3; the probability P (rtot > 6) is about 2/3. The mean of r1 is 6, and the mean of rtot is 11.

(2.110)

(2.111)

The probabilities of error in these representative cases are thus P (error | r = 000) =

f3 (1 − f )3 + f 3

(2.112)

and P (error | r = 001) = f.

(2.113)

Notice that while the average probability of error of R 3 is about 3f 2 , the probability (given r) that any particular bit is wrong is either about f 3 or f . The average error probability, using the sum rule, is X P (r)P (error | r) P (error) = r

f3 = 2[1/2(1 − f )3 + 1/2f 3 ] + 6[1/2f (1 − f )]f. (1 − f )3 + f 3

So P (error) = f 3 + 3f 2 (1 − f ). Solution to exercise 2.39 (p.40).

The entropy is 9.7 bits per word.

The first two terms are for the cases r = 000 and 111; the remaining 6 are for the other outcomes, which share the same probability of occurring and identical error probability, f .

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About Chapter 3 If you are eager to get on to information theory, data compression, and noisy channels, you can skip to Chapter 4. Data compression and data modelling are intimately connected, however, so you’ll probably want to come back to this chapter by the time you get to Chapter 6. Before reading Chapter 3, it might be good to look at the following exercises. . Exercise 3.1.[2, p.59] A die is selected at random from two twenty-faced dice on which the symbols 1–10 are written with nonuniform frequency as follows. Symbol

1

2

3

4

5

6

7

8

9

10

Number of faces of die A Number of faces of die B

6 3

4 3

3 2

2 2

1 2

1 2

1 2

1 2

1 1

0 1

The randomly chosen die is rolled 7 times, with the following outcomes: 5, 3, 9, 3, 8, 4, 7. What is the probability that the die is die A? . Exercise 3.2.[2, p.59] Assume that there is a third twenty-faced die, die C, on which the symbols 1–20 are written once each. As above, one of the three dice is selected at random and rolled 7 times, giving the outcomes: 3, 5, 4, 8, 3, 9, 7. What is the probability that the die is (a) die A, (b) die B, (c) die C? Exercise 3.3.[3, p.48] Inferring a decay constant Unstable particles are emitted from a source and decay at a distance x, a real number that has an exponential probability distribution with characteristic length λ. Decay events can be observed only if they occur in a window extending from x = 1 cm to x = 20 cm. N decays are observed at locations {x1 , . . . , xN }. What is λ?

* ** * *

*

*

* *

x . Exercise 3.4.[3, p.55] Forensic evidence Two people have left traces of their own blood at the scene of a crime. A suspect, Oliver, is tested and found to have type ‘O’ blood. The blood groups of the two traces are found to be of type ‘O’ (a common type in the local population, having frequency 60%) and of type ‘AB’ (a rare type, with frequency 1%). Do these data (type ‘O’ and ‘AB’ blood were found at scene) give evidence in favour of the proposition that Oliver was one of the two people present at the crime?

47

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3 More about Inference

It is not a controversial statement that Bayes’ theorem provides the correct language for describing the inference of a message communicated over a noisy channel, as we used it in Chapter 1 (p.6). But strangely, when it comes to other inference problems, the use of Bayes’ theorem is not so widespread.

3.1 A first inference problem When I was an undergraduate in Cambridge, I was privileged to receive supervisions from Steve Gull. Sitting at his desk in a dishevelled office in St. John’s College, I asked him how one ought to answer an old Tripos question (exercise 3.3): Unstable particles are emitted from a source and decay at a distance x, a real number that has an exponential probability distribution with characteristic length λ. Decay events can be observed only if they occur in a window extending from x = 1 cm to x = 20 cm. N decays are observed at locations {x 1 , . . . , xN }. What is λ?

* ** * *

*

*

* *

x I had scratched my head over this for some time. My education had provided me with a couple of approaches to solving such inference problems: constructing ‘estimators’ of the unknown parameters; or ‘fitting’ the model to the data, or to a processed version of the data. Since the mean of an unconstrained exponential distribution is λ, it seemed P reasonable to examine the sample mean x ¯ = n xn /N and see if an estimator ˆ could be obtained from it. It was evident that the estimator λ ˆ=x λ ¯ −1 would be appropriate for λ  20 cm, but not for cases where the truncation of the distribution at the right-hand side is significant; with a little ingenuity and the introduction of ad hoc bins, promising estimators for λ  20 cm could be constructed. But there was no obvious estimator that would work under all conditions. Nor could I find a satisfactory approach based on fitting the density P (x | λ) to a histogram derived from the data. I was stuck. What is the general solution to this problem and others like it? Is it always necessary, when confronted by a new inference problem, to grope in the dark for appropriate ‘estimators’ and worry about finding the ‘best’ estimator (whatever that means)? 48

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3.1: A first inference problem

49

0.25

Figure 3.1. The probability density P (x | λ) as a function of x.

P(x|lambda=2) P(x|lambda=5) P(x|lambda=10)

0.2 0.15 0.1 0.05 0 2

4

6

8

10

12

14

16

18

20

x

0.2

Figure 3.2. The probability density P (x | λ) as a function of λ, for three different values of x. When plotted this way round, the function is known as the likelihood of λ. The marks indicate the three values of λ, λ = 2, 5, 10, that were used in the preceding figure.

P(x=3|lambda) P(x=5|lambda) P(x=12|lambda)

0.15

0.1

0.05

0 1

10

100

λ

Steve wrote down the probability of one data point, given λ:  1 −x/λ e /Z(λ) 1 < x < 20 λ P (x | λ) = 0 otherwise where Z(λ) =

Z

20

dx 1

1 λ

  e−x/λ = e−1/λ − e−20/λ .

(3.1)

(3.2)

This seemed obvious enough. Then he wrote Bayes’ theorem: P (λ | {x1 , . . . , xN }) = ∝

P ({x} | λ)P (λ) P ({x})  P  1 N exp − x /λ P (λ). 1 n (λZ(λ))N

(3.3) (3.4)

Suddenly, the straightforward distribution P ({x 1 , . . . , xN } | λ), defining the probability of the data given the hypothesis λ, was being turned on its head so as to define the probability of a hypothesis given the data. A simple figure showed the probability of a single data point P (x | λ) as a familiar function of x, for different values of λ (figure 3.1). Each curve was an innocent exponential, normalized to have area 1. Plotting the same function as a function of λ for a fixed value of x, something remarkable happens: a peak emerges (figure 3.2). To help understand these two points of view of the one function, figure 3.3 shows a surface plot of P (x | λ) as a function of x and λ. For a dataset consisting of several points, e.g., the six points {x} N n=1 = {1.5, 2, 3, 4, 5, 12}, the likelihood function P ({x} | λ) is the product of the N functions of λ, P (xn | λ) (figure 3.4). 1.4e-06

3

2

1 100 10

1

x

1.5 1

2

λ

2.5

Figure 3.3. The probability density P (x | λ) as a function of x and λ. Figures 3.1 and 3.2 are vertical sections through this surface.

Figure 3.4. The likelihood function in the case of a six-point dataset, P ({x} = {1.5, 2, 3, 4, 5, 12} | λ), as a function of λ.

1.2e-06 1e-06 8e-07 6e-07 4e-07 2e-07 0 1

10

100

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50

3 — More about Inference Steve summarized Bayes’ theorem as embodying the fact that what you know about λ after the data arrive is what you knew before [P (λ)], and what the data told you [P ({x} | λ)].

Probabilities are used here to quantify degrees of belief. To nip possible confusion in the bud, it must be emphasized that the hypothesis λ that correctly describes the situation is not a stochastic variable, and the fact that the Bayesian uses a probability distribution P does not mean that he thinks of the world as stochastically changing its nature between the states described by the different hypotheses. He uses the notation of probabilities to represent his beliefs about the mutually exclusive micro-hypotheses (here, values of λ), of which only one is actually true. That probabilities can denote degrees of belief, given assumptions, seemed reasonable to me. The posterior probability distribution (3.4) represents the unique and complete solution to the problem. There is no need to invent ‘estimators’; nor do we need to invent criteria for comparing alternative estimators with each other. Whereas orthodox statisticians offer twenty ways of solving a problem, and another twenty different criteria for deciding which of these solutions is the best, Bayesian statistics only offers one answer to a well-posed problem.

Assumptions in inference Our inference is conditional on our assumptions [for example, the prior P (λ)]. Critics view such priors as a difficulty because they are ‘subjective’, but I don’t see how it could be otherwise. How can one perform inference without making assumptions? I believe that it is of great value that Bayesian methods force one to make these tacit assumptions explicit. First, once assumptions are made, the inferences are objective and unique, reproducible with complete agreement by anyone who has the same information and makes the same assumptions. For example, given the assumptions listed above, H, and the data D, everyone will agree about the posterior probability of the decay length λ: P (λ | D, H) =

P (D | λ, H)P (λ | H) . P (D | H)

(3.5)

Second, when the assumptions are explicit, they are easier to criticize, and easier to modify – indeed, we can quantify the sensitivity of our inferences to the details of the assumptions. For example, we can note from the likelihood curves in figure 3.2 that in the case of a single data point at x = 5, the likelihood function is less strongly peaked than in the case x = 3; the details of the prior P (λ) become increasingly important as the sample mean x ¯ gets closer to the middle of the window, 10.5. In the case x = 12, the likelihood function doesn’t have a peak at all – such data merely rule out small values of λ, and don’t give any information about the relative probabilities of large values of λ. So in this case, the details of the prior at the small–λ end of things are not important, but at the large–λ end, the prior is important. Third, when we are not sure which of various alternative assumptions is the most appropriate for a problem, we can treat this question as another inference task. Thus, given data D, we can compare alternative assumptions H using Bayes’ theorem: P (H | D, I) =

P (D | H, I)P (H | I) , P (D | I)

(3.6)

If you have any difficulty understanding this chapter I recommend ensuring you are happy with exercises 3.1 and 3.2 (p.47) then noting their similarity to exercise 3.3.

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3.2: The bent coin

51

where I denotes the highest assumptions, which we are not questioning. Fourth, we can take into account our uncertainty regarding such assumptions when we make subsequent predictions. Rather than choosing one particular assumption H∗ , and working out our predictions about some quantity t, P (t | D, H∗ , I), we obtain predictions that take into account our uncertainty about H by using the sum rule: X P (t | D, I) = P (t | D, H, I)P (H | D, I). (3.7) H

This is another contrast with orthodox statistics, in which it is conventional to ‘test’ a default model, and then, if the test ‘accepts the model’ at some ‘significance level’, to use exclusively that model to make predictions. Steve thus persuaded me that probability theory reaches parts that ad hoc methods cannot reach. Let’s look at a few more examples of simple inference problems.

3.2 The bent coin A bent coin is tossed F times; we observe a sequence s of heads and tails (which we’ll denote by the symbols a and b). We wish to know the bias of the coin, and predict the probability that the next toss will result in a head. We first encountered this task in example 2.7 (p.30), and we will encounter it again in Chapter 6, when we discuss adaptive data compression. It is also the original inference problem studied by Thomas Bayes in his essay published in 1763. As in exercise 2.8 (p.30), we will assume a uniform prior distribution and obtain a posterior distribution by multiplying by the likelihood. A critic might object, ‘where did this prior come from?’ I will not claim that the uniform prior is in any way fundamental; indeed we’ll give examples of nonuniform priors later. The prior is a subjective assumption. One of the themes of this book is: you can’t do inference – or data compression – without making assumptions. We give the name H1 to our assumptions. [We’ll be introducing an alternative set of assumptions in a moment.] The probability, given p a , that F tosses result in a sequence s that contains {F a , Fb } counts of the two outcomes is (3.8) P (s | pa , F, H1 ) = pFa a (1 − pa )Fb . [For example, P (s = aaba | pa , F = 4, H1 ) = pa pa (1 − pa )pa .] Our first model assumes a uniform prior distribution for p a , P (pa | H1 ) = 1,

pa ∈ [0, 1]

(3.9)

and pb ≡ 1 − pa . Inferring unknown parameters Given a string of length F of which Fa are as and Fb are bs, we are interested in (a) inferring what pa might be; (b) predicting whether the next character is

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52

3 — More about Inference

an a or a b. [Predictions are always expressed as probabilities. So ‘predicting whether the next character is an a’ is the same as computing the probability that the next character is an a.] Assuming H1 to be true, the posterior probability of p a , given a string s of length F that has counts {Fa , Fb }, is, by Bayes’ theorem, P (pa | s, F, H1 ) =

P (s | pa , F, H1 )P (pa | H1 ) . P (s | F, H1 )

(3.10)

pFa a (1 − pa )Fb . P (s | F, H1 )

(3.11)

The factor P (s | pa , F, H1 ), which, as a function of pa , is known as the likelihood function, was given in equation (3.8); the prior P (p a | H1 ) was given in equation (3.9). Our inference of pa is thus: P (pa | s, F, H1 ) =

The normalizing constant is given by the beta integral Z 1 Γ(Fa + 1)Γ(Fb + 1) Fa !Fb ! dpa pFa a (1 − pa )Fb = = . P (s | F, H1 ) = Γ(F + F + 2) (F + Fb + 1)! a b a 0 (3.12) Exercise 3.5.[2, p.59] Sketch the posterior probability P (p a | s = aba, F = 3). What is the most probable value of pa (i.e., the value that maximizes the posterior probability density)? What is the mean value of p a under this distribution? Answer the same P (pa | s = bbb, F = 3).

questions

for

the

posterior

probability

From inferences to predictions Our prediction about the next toss, the probability that the next toss is an a, is obtained by integrating over pa . This has the effect of taking into account our uncertainty about pa when making predictions. By the sum rule, Z P (a | s, F ) = dpa P (a | pa )P (pa | s, F ). (3.13) The probability of an a given pa is simply pa , so Z pFa (1 − pa )Fb P (a | s, F ) = dpa pa a P (s | F ) Z pFa a +1 (1 − pa )Fb = dpa P (s | F )     (Fa + 1)! Fb ! Fa ! F b ! Fa + 1 = = , (Fa + Fb + 2)! (Fa + Fb + 1)! Fa + F b + 2

(3.14) (3.15) (3.16)

which is known as Laplace’s rule.

3.3 The bent coin and model comparison Imagine that a scientist introduces another theory for our data. He asserts that the source is not really a bent coin but is really a perfectly formed die with one face painted heads (‘a’) and the other five painted tails (‘b’). Thus the parameter pa , which in the original model, H1 , could take any value between 0 and 1, is according to the new hypothesis, H 0 , not a free parameter at all; rather, it is equal to 1/6. [This hypothesis is termed H 0 so that the suffix of each model indicates its number of free parameters.] How can we compare these two models in the light of data? We wish to infer how probable H1 is relative to H0 .

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3.3: The bent coin and model comparison

53

Model comparison as inference In order to perform model comparison, we write down Bayes’ theorem again, but this time with a different argument on the left-hand side. We wish to know how probable H1 is given the data. By Bayes’ theorem, P (H1 | s, F ) =

P (s | F, H1 )P (H1 ) . P (s | F )

(3.17)

Similarly, the posterior probability of H 0 is P (H0 | s, F ) =

P (s | F, H0 )P (H0 ) . P (s | F )

(3.18)

The normalizing constant in both cases is P (s | F ), which is the total probability of getting the observed data. If H 1 and H0 are the only models under consideration, this probability is given by the sum rule: P (s | F ) = P (s | F, H1 )P (H1 ) + P (s | F, H0 )P (H0 ).

(3.19)

To evaluate the posterior probabilities of the hypotheses we need to assign values to the prior probabilities P (H 1 ) and P (H0 ); in this case, we might set these to 1/2 each. And we need to evaluate the data-dependent terms P (s | F, H1 ) and P (s | F, H0 ). We can give names to these quantities. The quantity P (s | F, H1 ) is a measure of how much the data favour H 1 , and we call it the evidence for model H1 . We already encountered this quantity in equation (3.10) where it appeared as the normalizing constant of the first inference we made – the inference of p a given the data. How model comparison works: The evidence for a model is usually the normalizing constant of an earlier Bayesian inference. We evaluated the normalizing constant for model H 1 in (3.12). The evidence for model H0 is very simple because this model has no parameters to infer. Defining p0 to be 1/6, we have P (s | F, H0 ) = pF0 a (1 − p0 )Fb .

(3.20)

Thus the posterior probability ratio of model H 1 to model H0 is P (H1 | s, F ) P (H0 | s, F )

= =

P (s | F, H1 )P (H1 ) P (s | F, H0 )P (H0 )  Fa !Fb ! pF0 a (1 − p0 )Fb . (Fa + Fb + 1)!

(3.21) (3.22)

Some values of this posterior probability ratio are illustrated in table 3.5. The first five lines illustrate that some outcomes favour one model, and some favour the other. No outcome is completely incompatible with either model. With small amounts of data (six tosses, say) it is typically not the case that one of the two models is overwhelmingly more probable than the other. But with more data, the evidence against H0 given by any data set with the ratio F a : Fb differing from 1: 5 mounts up. You can’t predict in advance how much data are needed to be pretty sure which theory is true. It depends what p a is. The simpler model, H0 , since it has no adjustable parameters, is able to lose out by the biggest margin. The odds may be hundreds to one against it. The more complex model can never lose out by a large margin; there’s no data set that is actually unlikely given model H 1 .

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54

3 — More about Inference

F

Data (Fa , Fb )

6 6 6 6 6

(5, 1) (3, 3) (2, 4) (1, 5) (0, 6)

20 20 20

(10, 10) (3, 17) (0, 20)

222.2 2.67 0.71 0.356 0.427

pa = 1/6

0

50

1000/1 100/1 10/1 1/1 1/10 1/100 100 150 200

50

1000/1 100/1 10/1 1/1 1/10 1/100 100 150 200

50

1000/1 100/1 10/1 1/1 1/10 1/100 100 150 200

8 6 4 2 0 -2 -4 0 8 6 4 2 0 -2 -4 0

= 1/1.4 = 1/2.8 = 1/2.3

96.5 0.2 1.83

H0 is true 8 6 4 2 0 -2 -4

Table 3.5. Outcome of model comparison between models H1 and H0 for the ‘bent coin’. Model H0 states that pa = 1/6, pb = 5/6.

P (H1 | s, F ) P (H0 | s, F )

= 1/5

H1 is true pa = 0.25

8 6 4 2 0 -2 -4 0

50

1000/1 100/1 10/1 1/1 1/10 1/100 100 150 200

50

1000/1 100/1 10/1 1/1 1/10 1/100 100 150 200

50

1000/1 100/1 10/1 1/1 1/10 1/100 100 150 200

8 6 4 2 0 -2 -4 0 8 6 4 2 0 -2 -4 0

pa = 0.5

8 6 4 2 0 -2 -4 0

50

1000/1 100/1 10/1 1/1 1/10 1/100 100 150 200

50

1000/1 100/1 10/1 1/1 1/10 1/100 100 150 200

50

1000/1 100/1 10/1 1/1 1/10 1/100 100 150 200

8 6 4 2 0 -2 -4 0 8 6 4 2 0 -2 -4 0

. Exercise 3.6.[2 ] Show that after F tosses have taken place, the biggest value that the log evidence ratio log

P (s | F, H1 ) P (s | F, H0 )

(3.23)

can have scales linearly with F if H1 is more probable, but the log evidence in favour of H0 can grow at most as log F . . Exercise 3.7.[3, p.60] Putting your sampling theory hat on, assuming F a has not yet been measured, compute a plausible range that the log evidence ratio might lie in, as a function of F and the true value of p a , and sketch it as a function of F for pa = p0 = 1/6, pa = 0.25, and pa = 1/2. [Hint: sketch the log evidence as a function of the random variable F a and work out the mean and standard deviation of F a .]

Typical behaviour of the evidence Figure 3.6 shows the log evidence ratio as a function of the number of tosses, F , in a number of simulated experiments. In the left-hand experiments, H 0 was true. In the right-hand ones, H1 was true, and the value of pa was either 0.25 or 0.5. We will discuss model comparison more in a later chapter.

Figure 3.6. Typical behaviour of the evidence in favour of H1 as bent coin tosses accumulate under three different conditions (columns 1, 2, 3). Horizontal axis is the number of tosses, F . The vertical axis on the left is (s | F, H1 ) ln P ; the right-hand P (s | F, H0 ) vertical axis shows the values of P (s | F, H1 ) . P (s | F, H0 ) The three rows show independent simulated experiments. (See also figure 3.8, p.60.)

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3.4: An example of legal evidence

55

3.4 An example of legal evidence The following example illustrates that there is more to Bayesian inference than the priors. Two people have left traces of their own blood at the scene of a crime. A suspect, Oliver, is tested and found to have type ‘O’ blood. The blood groups of the two traces are found to be of type ‘O’ (a common type in the local population, having frequency 60%) and of type ‘AB’ (a rare type, with frequency 1%). Do these data (type ‘O’ and ‘AB’ blood were found at scene) give evidence in favour of the proposition that Oliver was one of the two people present at the crime? A careless lawyer might claim that the fact that the suspect’s blood type was found at the scene is positive evidence for the theory that he was present. But this is not so. Denote the proposition ‘the suspect and one unknown person were present’ ¯ states ‘two unknown people from the population were by S. The alternative, S, present’. The prior in this problem is the prior probability ratio between the ¯ This quantity is important to the final verdict and propositions S and S. would be based on all other available information in the case. Our task here is just to evaluate the contribution made by the data D, that is, the likelihood ¯ H). In my view, a jury’s task should generally be to ratio, P (D | S, H)/P (D | S, multiply together carefully evaluated likelihood ratios from each independent piece of admissible evidence with an equally carefully reasoned prior probability. [This view is shared by many statisticians but learned British appeal judges recently disagreed and actually overturned the verdict of a trial because the jurors had been taught to use Bayes’ theorem to handle complicated DNA evidence.] The probability of the data given S is the probability that one unknown person drawn from the population has blood type AB: P (D | S, H) = pAB

(3.24)

(since given S, we already know that one trace will be of type O). The probability of the data given S¯ is the probability that two unknown people drawn from the population have types O and AB: ¯ H) = 2 pO pAB . P (D | S,

(3.25)

In these equations H denotes the assumptions that two people were present and left blood there, and that the probability distribution of the blood groups of unknown people in an explanation is the same as the population frequencies. Dividing, we obtain the likelihood ratio: 1 P (D | S, H) 1 = = 0.83. = ¯ 2pO 2 × 0.6 P (D | S, H)

(3.26)

Thus the data in fact provide weak evidence against the supposition that Oliver was present. This result may be found surprising, so let us examine it from various points of view. First consider the case of another suspect, Alberto, who has type AB. Intuitively, the data do provide evidence in favour of the theory S 0

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3 — More about Inference

¯ And indeed that this suspect was present, relative to the null hypothesis S. the likelihood ratio in this case is: 1 P (D | S 0 , H) ¯ H) = 2 pAB = 50. P (D | S,

(3.27)

Now let us change the situation slightly; imagine that 99% of people are of blood type O, and the rest are of type AB. Only these two blood types exist in the population. The data at the scene are the same as before. Consider again how these data influence our beliefs about Oliver, a suspect of type O, and Alberto, a suspect of type AB. Intuitively, we still believe that the presence of the rare AB blood provides positive evidence that Alberto was there. But does the fact that type O blood was detected at the scene favour the hypothesis that Oliver was present? If this were the case, that would mean that regardless of who the suspect is, the data make it more probable they were present; everyone in the population would be under greater suspicion, which would be absurd. The data may be compatible with any suspect of either blood type being present, but if they provide evidence for some theories, they must also provide evidence against other theories. Here is another way of thinking about this: imagine that instead of two people’s blood stains there are ten, and that in the entire local population of one hundred, there are ninety type O suspects and ten type AB suspects. Consider a particular type O suspect, Oliver: without any other information, and before the blood test results come in, there is a one in 10 chance that he was at the scene, since we know that 10 out of the 100 suspects were present. We now get the results of blood tests, and find that nine of the ten stains are of type AB, and one of the stains is of type O. Does this make it more likely that Oliver was there? No, there is now only a one in ninety chance that he was there, since we know that only one person present was of type O. Maybe the intuition is aided finally by writing down the formulae for the general case where nO blood stains of individuals of type O are found, and nAB of type AB, a total of N individuals in all, and unknown people come from a large population with fractions p O , pAB . (There may be other blood types too.) The task is to evaluate the likelihood ratio for the two hypotheses: S, ‘the type O suspect (Oliver) and N −1 unknown others left N stains’; and ¯ ‘N unknowns left N stains’. The probability of the data under hypothesis S, ¯ S is just the probability of getting n O , nAB individuals of the two types when N individuals are drawn at random from the population: ¯ = P (nO , nAB | S)

N! pnO pnAB . nO ! nAB ! O AB

(3.28)

In the case of hypothesis S, we need the distribution of the N −1 other individuals: (N − 1)! AB P (nO , nAB | S) = . (3.29) pnO −1 pnAB (nO − 1)! nAB ! O The likelihood ratio is:

P (nO , nAB | S) nO /N ¯ = pO . P (nO , nAB | S)

(3.30)

This is an instructive result. The likelihood ratio, i.e. the contribution of these data to the question of whether Oliver was present, depends simply on a comparison of the frequency of his blood type in the observed data with the background frequency in the population. There is no dependence on the counts of the other types found at the scene, or their frequencies in the population.

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3.5: Exercises If there are more type O stains than the average number expected under ¯ then the data give evidence in favour of the presence of Oliver. hypothesis S, Conversely, if there are fewer type O stains than the expected number under ¯ then the data reduce the probability of the hypothesis that he was there. S, In the special case nO /N = pO , the data contribute no evidence either way, regardless of the fact that the data are compatible with the hypothesis S.

3.5 Exercises Exercise 3.8.[2, p.60] The three doors, normal rules. On a game show, a contestant is told the rules as follows: There are three doors, labelled 1, 2, 3. A single prize has been hidden behind one of them. You get to select one door. Initially your chosen door will not be opened. Instead, the gameshow host will open one of the other two doors, and he will do so in such a way as not to reveal the prize. For example, if you first choose door 1, he will then open one of doors 2 and 3, and it is guaranteed that he will choose which one to open so that the prize will not be revealed. At this point, you will be given a fresh choice of door: you can either stick with your first choice, or you can switch to the other closed door. All the doors will then be opened and you will receive whatever is behind your final choice of door. Imagine that the contestant chooses door 1 first; then the gameshow host opens door 3, revealing nothing behind the door, as promised. Should the contestant (a) stick with door 1, or (b) switch to door 2, or (c) does it make no difference? Exercise 3.9.[2, p.61] The three doors, earthquake scenario. Imagine that the game happens again and just as the gameshow host is about to open one of the doors a violent earthquake rattles the building and one of the three doors flies open. It happens to be door 3, and it happens not to have the prize behind it. The contestant had initially chosen door 1. Repositioning his toup´ee, the host suggests, ‘OK, since you chose door 1 initially, door 3 is a valid door for me to open, according to the rules of the game; I’ll let door 3 stay open. Let’s carry on as if nothing happened.’ Should the contestant stick with door 1, or switch to door 2, or does it make no difference? Assume that the prize was placed randomly, that the gameshow host does not know where it is, and that the door flew open because its latch was broken by the earthquake. [A similar alternative scenario is a gameshow whose confused host forgets the rules, and where the prize is, and opens one of the unchosen doors at random. He opens door 3, and the prize is not revealed. Should the contestant choose what’s behind door 1 or door 2? Does the optimal decision for the contestant depend on the contestant’s beliefs about whether the gameshow host is confused or not?] . Exercise 3.10.[2 ] Another example in which the emphasis is not on priors. You visit a family whose three children are all at the local school. You don’t

57

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58

3 — More about Inference know anything about the sexes of the children. While walking clumsily round the home, you stumble through one of the three unlabelled bedroom doors that you know belong, one each, to the three children, and find that the bedroom contains girlie stuff in sufficient quantities to convince you that the child who lives in that bedroom is a girl. Later, you sneak a look at a letter addressed to the parents, which reads ‘From the Headmaster: we are sending this letter to all parents who have male children at the school to inform them about the following boyish matters. . . ’. These two sources of evidence establish that at least one of the three children is a girl, and that at least one of the children is a boy. What are the probabilities that there are (a) two girls and one boy; (b) two boys and one girl?

. Exercise 3.11.[2, p.61] Mrs S is found stabbed in her family garden. Mr S behaves strangely after her death and is considered as a suspect. On investigation of police and social records it is found that Mr S had beaten up his wife on at least nine previous occasions. The prosecution advances this data as evidence in favour of the hypothesis that Mr S is guilty of the murder. ‘Ah no,’ says Mr S’s highly paid lawyer, ‘statistically, only one in a thousand wife-beaters actually goes on to murder his wife. 1 So the wife-beating is not strong evidence at all. In fact, given the wife-beating evidence alone, it’s extremely unlikely that he would be the murderer of his wife – only a 1/1000 chance. You should therefore find him innocent.’ Is the lawyer right to imply that the history of wife-beating does not point to Mr S’s being the murderer? Or is the lawyer a slimy trickster? If the latter, what is wrong with his argument? [Having received an indignant letter from a lawyer about the preceding paragraph, I’d like to add an extra inference exercise at this point: Does my suggestion that Mr. S.’s lawyer may have been a slimy trickster imply that I believe all lawyers are slimy tricksters? (Answer: No.)] . Exercise 3.12.[2 ] A bag contains one counter, known to be either white or black. A white counter is put in, the bag is shaken, and a counter is drawn out, which proves to be white. What is now the chance of drawing a white counter? [Notice that the state of the bag, after the operations, is exactly identical to its state before.] . Exercise 3.13.[2, p.62] You move into a new house; the phone is connected, and you’re pretty sure that the phone number is 740511, but not as sure as you would like to be. As an experiment, you pick up the phone and dial 740511; you obtain a ‘busy’ signal. Are you now more sure of your phone number? If so, how much? . Exercise 3.14.[1 ] In a game, two coins are tossed. If either of the coins comes up heads, you have won a prize. To claim the prize, you must point to one of your coins that is a head and say ‘look, that coin’s a head, I’ve won’. You watch Fred play the game. He tosses the two coins, and he 1 In the U.S.A., it is estimated that 2 million women are abused each year by their partners. In 1994, 4739 women were victims of homicide; of those, 1326 women (28%) were slain by husbands and boyfriends. (Sources: http://www.umn.edu/mincava/papers/factoid.htm, http://www.gunfree.inter.net/vpc/womenfs.htm)

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3.6: Solutions

59

points to a coin and says ‘look, that coin’s a head, I’ve won’. What is the probability that the other coin is a head? . Exercise 3.15.[2, p.63] A statistical statement appeared in The Guardian on Friday January 4, 2002: When spun on edge 250 times, a Belgian one-euro coin came up heads 140 times and tails 110. ‘It looks very suspicious to me’, said Barry Blight, a statistics lecturer at the London School of Economics. ‘If the coin were unbiased the chance of getting a result as extreme as that would be less than 7%’. But do these data give evidence that the coin is biased rather than fair? [Hint: see equation (3.22).]

3.6 Solutions Solution to exercise 3.1 (p.47). probabilities,

Let the data be D. Assuming equal prior

1313121 9 P (A | D) = = P (B | D) 2212222 32

(3.31)

and P (A | D) = 9/41. Solution to exercise 3.2 (p.47). pothesis is:

The probability of the data given each hy-

P (D | A) =

(3.32)

3 1 2 1 3 1 1 18 = 7; 20 20 20 20 20 20 20 20 64 2 2 2 2 2 1 2 = 7; P (D | B) = 20 20 20 20 20 20 20 20 1 1 1 1 1 1 1 1 P (D | C) = = 7. 20 20 20 20 20 20 20 20

(3.33) (3.34)

So P (A | D) =

18 18 = ; 18 + 64 + 1 83

P (B | D) =

64 ; 83

P (C | D) =

1 . 83 (3.35) Figure 3.7. Posterior probability for the bias pa of a bent coin given two different data sets.

0 0.2 0.4 0.6 0.8 1 (a) P (pa | s = aba, F = 3) ∝ p2a (1 − pa )

(b)

0

0.2

0.4

0.6

0.8

1

P (pa | s = bbb, F = 3) ∝ (1 − pa )3

Solution to exercise 3.5 (p.52). (a) P (pa | s = aba, F = 3) ∝ p2a (1 − pa ). The most probable value of pa (i.e., the value that maximizes the posterior probability density) is 2/3. The mean value of pa is 3/5. See figure 3.7a.

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3 — More about Inference

(b) P (pa | s = bbb, F = 3) ∝ (1 − pa )3 . The most probable value of pa (i.e., the value that maximizes the posterior probability density) is 0. The mean value of pa is 1/5. See figure 3.7b. H0 is true

H1 is true

pa = 1/6

8 6 4 2 0 -2 -4 0

50

1000/1 100/1 10/1 1/1 1/10 1/100 100 150 200

pa = 0.25

8 6 4 2 0 -2 -4 0

50

1000/1 100/1 10/1 1/1 1/10 1/100 100 150 200

pa = 0.5

8 6 4 2 0 -2 -4 0

50

1000/1 100/1 10/1 1/1 1/10 1/100 100 150 200

Solution to exercise 3.7 (p.54). The curves in figure 3.8 were found by finding the mean and standard deviation of F a , then setting Fa to the mean ± two standard deviations to get a 95% plausible range for F a , and computing the three corresponding values of the log evidence ratio. Solution to exercise 3.8 (p.57). Let H i denote the hypothesis that the prize is behind door i. We make the following assumptions: the three hypotheses H 1 , H2 and H3 are equiprobable a priori, i.e., 1 P (H1 ) = P (H2 ) = P (H3 ) = . 3

(3.36)

The datum we receive, after choosing door 1, is one of D = 3 and D = 2 (meaning door 3 or 2 is opened, respectively). We assume that these two possible outcomes have the following probabilities. If the prize is behind door 1 then the host has a free choice; in this case we assume that the host selects at random between D = 2 and D = 3. Otherwise the choice of the host is forced and the probabilities are 0 and 1. P (D = 2 | H1 ) = 1/2 P (D = 2 | H2 ) = 0 P (D = 2 | H3 ) = 1 P (D = 3 | H1 ) = 1/2 P (D = 3 | H2 ) = 1 P (D = 3 | H3 ) = 0

(3.37)

Now, using Bayes’ theorem, we evaluate the posterior probabilities of the hypotheses: P (D = 3 | Hi )P (Hi ) P (Hi | D = 3) = (3.38) P (D = 3) P (H1 | D = 3) = (1/2)(1/3) P (D=3)

P (H2 | D = 3) = P(1)(1/3) (D=3)

P (H3 | D = 3) = P(0)(1/3) (D=3) (3.39) The denominator P (D = 3) is (1/2) because it is the normalizing constant for this posterior distribution. So P (H1 | D = 3) = 1/3 P (H2 | D = 3) = 2/3 P (H3 | D = 3) = 0. (3.40) So the contestant should switch to door 2 in order to have the biggest chance of getting the prize. Many people find this outcome surprising. There are two ways to make it more intuitive. One is to play the game thirty times with a friend and keep track of the frequency with which switching gets the prize. Alternatively, you can perform a thought experiment in which the game is played with a million doors. The rules are now that the contestant chooses one door, then the game

Figure 3.8. Range of plausible values of the log evidence in favour of H1 as a function of F . The vertical axis on the left is (s | F,H1 ) log P ; the right-hand P (s | F,H0 ) vertical axis shows the values of P (s | F,H1 ) . P (s | F,H0 ) The solid line shows the log evidence if the random variable Fa takes on its mean value, Fa = pa F . The dotted lines show (approximately) the log evidence if Fa is at its 2.5th or 97.5th percentile. (See also figure 3.6, p.54.)

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3.6: Solutions

61

show host opens 999,998 doors in such a way as not to reveal the prize, leaving the contestant’s selected door and one other door closed. The contestant may now stick or switch. Imagine the contestant confronted by a million doors, of which doors 1 and 234,598 have not been opened, door 1 having been the contestant’s initial guess. Where do you think the prize is?

P (D|H1 )(1/3) P (D)

P (H1 |D) = = 1/2

P (D|H2 )(1/3) P (D)

P (H2 |D) = = 1/2

P (D|H3 )(1/3) P (D)

P (H3 |D) = = 0.

(3.41) The two possible hypotheses are now equally likely. If we assume that the host knows where the prize is and might be acting deceptively, then the answer might be further modified, because we have to view the host’s words as part of the data. Confused? It’s well worth making sure you understand these two gameshow problems. Don’t worry, I slipped up on the second problem, the first time I met it. There is a general rule which helps immensely when you have a confusing probability problem: Always write down the probability of everything. (Steve Gull) From this joint probability, any desired inference can be mechanically obtained (figure 3.9). Solution to exercise 3.11 (p.58). The statistic quoted by the lawyer indicates the probability that a randomly selected wife-beater will also murder his wife. The probability that the husband was the murderer, given that the wife has been murdered, is a completely different quantity.

Where the prize is door door door 1 2 3 none

Which doors opened by earthquake

Solution to exercise 3.9 (p.57). If door 3 is opened by an earthquake, the inference comes out differently – even though visually the scene looks the same. The nature of the data, and the probability of the data, are both now different. The possible data outcomes are, firstly, that any number of the doors might have opened. We could label the eight possible outcomes d = (0, 0, 0), (0, 0, 1), (0, 1, 0), (1, 0, 0), (0, 1, 1), . . . , (1, 1, 1). Secondly, it might be that the prize is visible after the earthquake has opened one or more doors. So the data D consists of the value of d, and a statement of whether the prize was revealed. It is hard to say what the probabilities of these outcomes are, since they depend on our beliefs about the reliability of the door latches and the properties of earthquakes, but it is possible to extract the desired posterior probability without naming the values of P (d | H i ) for each d. All that matters are the relative values of the quantities P (D | H 1 ), P (D | H2 ), P (D | H3 ), for the value of D that actually occurred. [This is the likelihood principle, which we met in section 2.3.] The value of D that actually occurred is ‘d = (0, 0, 1), and no prize visible’. First, it is clear that P (D | H 3 ) = 0, since the datum that no prize is visible is incompatible with H 3 . Now, assuming that the contestant selected door 1, how does the probability P (D | H 1 ) compare with P (D | H2 )? Assuming that earthquakes are not sensitive to decisions of game show contestants, these two quantities have to be equal, by symmetry. We don’t know how likely it is that door 3 falls off its hinges, but however likely it is, it’s just as likely to do so whether the prize is behind door 1 or door 2. So, if P (D | H1 ) and P (D | H2 ) are equal, we obtain:

pnone pnone pnone 3 3 3

1 2 3

p3 3

p3 3

p3 3

1,2 1,3 2,3 1,2,3

p1,2,3 p1,2,3 p1,2,3 3 3 3

Figure 3.9. The probability of everything, for the second three-door problem, assuming an earthquake has just occurred. Here, p3 is the probability that door 3 alone is opened by an earthquake.

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62

3 — More about Inference

To deduce the latter, we need to make further assumptions about the probability that the wife is murdered by someone else. If she lives in a neighbourhood with frequent random murders, then this probability is large and the posterior probability that the husband did it (in the absence of other evidence) may not be very large. But in more peaceful regions, it may well be that the most likely person to have murdered you, if you are found murdered, is one of your closest relatives. Let’s work out some illustrative numbers with the help of the statistics on page 58. Let m = 1 denote the proposition that a woman has been murdered; h = 1, the proposition that the husband did it; and b = 1, the proposition that he beat her in the year preceding the murder. The statement ‘someone else did it’ is denoted by h = 0. We need to define P (h | m = 1), P (b | h = 1, m = 1), and P (b = 1 | h = 0, m = 1) in order to compute the posterior probability P (h = 1 | b = 1, m = 1). From the statistics, we can read out P (h = 1 | m = 1) = 0.28. And if two million women out of 100 million are beaten, then P (b = 1 | h = 0, m = 1) = 0.02. Finally, we need a value for P (b | h = 1, m = 1): if a man murders his wife, how likely is it that this is the first time he laid a finger on her? I expect it’s pretty unlikely; so maybe P (b = 1 | h = 1, m = 1) is 0.9 or larger. By Bayes’ theorem, then, P (h = 1 | b = 1, m = 1) =

.9 × .28 ' 95%. .9 × .28 + .02 × .72

(3.42)

One way to make obvious the sliminess of the lawyer on p.58 is to construct arguments, with the same logical structure as his, that are clearly wrong. For example, the lawyer could say ‘Not only was Mrs. S murdered, she was murdered between 4.02pm and 4.03pm. Statistically, only one in a million wife-beaters actually goes on to murder his wife between 4.02pm and 4.03pm. So the wife-beating is not strong evidence at all. In fact, given the wife-beating evidence alone, it’s extremely unlikely that he would murder his wife in this way – only a 1/1,000,000 chance.’ Solution to exercise 3.13 (p.58). There are two hypotheses. H 0 : your number is 740511; H1 : it is another number. The data, D, are ‘when I dialed 740511, I got a busy signal’. What is the probability of D, given each hypothesis? If your number is 740511, then we expect a busy signal with certainty: P (D | H0 ) = 1. On the other hand, if H1 is true, then the probability that the number dialled returns a busy signal is smaller than 1, since various other outcomes were also possible (a ringing tone, or a number-unobtainable signal, for example). The value of this probability P (D | H1 ) will depend on the probability α that a random phone number similar to your own phone number would be a valid phone number, and on the probability β that you get a busy signal when you dial a valid phone number. I estimate from the size of my phone book that Cambridge has about 75 000 valid phone numbers, all of length six digits. The probability that a random six-digit number is valid is therefore about 75 000/10 6 = 0.075. If we exclude numbers beginning with 0, 1, and 9 from the random choice, the probability α is about 75 000/700 000 ' 0.1. If we assume that telephone numbers are clustered then a misremembered number might be more likely to be valid than a randomly chosen number; so the probability, α, that our guessed number would be valid, assuming H 1 is true, might be bigger than

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3.6: Solutions

63

0.1. Anyway, α must be somewhere between 0.1 and 1. We can carry forward this uncertainty in the probability and see how much it matters at the end. The probability β that you get a busy signal when you dial a valid phone number is equal to the fraction of phones you think are in use or off-the-hook when you make your tentative call. This fraction varies from town to town and with the time of day. In Cambridge, during the day, I would guess that about 1% of phones are in use. At 4am, maybe 0.1%, or fewer. The probability P (D | H1 ) is the product of α and β, that is, about 0.1 × 0.01 = 10−3 . According to our estimates, there’s about a one-in-a-thousand chance of getting a busy signal when you dial a random number; or one-in-ahundred, if valid numbers are strongly clustered; or one-in-10 4 , if you dial in the wee hours. How do the data affect your beliefs about your phone number? The posterior probability ratio is the likelihood ratio times the prior probability ratio: P (D | H0 ) P (H0 ) P (H0 | D) = . P (H1 | D) P (D | H1 ) P (H1 )

(3.43)

The likelihood ratio is about 100-to-1 or 1000-to-1, so the posterior probability ratio is swung by a factor of 100 or 1000 in favour of H 0 . If the prior probability of H0 was 0.5 then the posterior probability is P (H0 | D) =

1 1+

P (H1 | D) P (H0 | D)

' 0.99 or 0.999.

(3.44)

Solution to exercise 3.15 (p.59). We compare the models H 0 – the coin is fair – and H1 – the coin is biased, with the prior on its bias set to the uniform distribution P (p|H1 ) = 1. [The use of a uniform prior seems reasonable to me, since I know that some coins, such as American pennies, have severe biases when spun on edge; so the situations p = 0.01 or p = 0.1 or p = 0.95 would not surprise me.] When I mention H0 – the coin is fair – a pedant would say, ‘how absurd to even consider that the coin is fair – any coin is surely biased to some extent’. And of course I would agree. So will pedants kindly understand H0 as meaning ‘the coin is fair to within one part in a thousand, i.e., p ∈ 0.5 ± 0.001’.

The likelihood ratio is: 140!110!

P (D|H1 ) = 251! = 0.48. P (D|H0 ) 1/2250

(3.45)

Thus the data give scarcely any evidence either way; in fact they give weak evidence (two to one) in favour of H0 ! ‘No, no’, objects the believer in bias, ‘your silly uniform prior doesn’t represent my prior beliefs about the bias of biased coins – I was expecting only a small bias’. To be as generous as possible to the H 1 , let’s see how well it could fare if the prior were presciently set. Let us allow a prior of the form P (p|H1 , α) =

1 α−1 p (1 − p)α−1 , Z(α)

where Z(α) = Γ(α)2 /Γ(2α)

(3.46)

(a Beta distribution, with the original uniform prior reproduced by setting α = 1). By tweaking α, the likelihood ratio for H 1 over H0 , Γ(140+α) Γ(110+α) Γ(2α)2250 P (D|H1 , α) = , P (D|H0 ) Γ(250+2α) Γ(α)2

(3.47)

0.05

H0 H1

0.04 0.03 0.02

140

0.01 0 0

50

100

150

200

250

Figure 3.10. The probability distribution of the number of heads given the two hypotheses, that the coin is fair, and that it is biased, with the prior distribution of the bias being uniform. The outcome (D = 140 heads) gives weak evidence in favour of H0 , the hypothesis that the coin is fair.

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64

3 — More about Inference

can be increased a little. It is shown for several values of α in figure 3.11. Even the most favourable choice of α (α ' 50) can yield a likelihood ratio of only two to one in favour of H1 . In conclusion, the data are not ‘very suspicious’. They can be construed as giving at most two-to-one evidence in favour of one or other of the two hypotheses. Are these wimpy likelihood ratios the fault of over-restrictive priors? Is there any way of producing a ‘very suspicious’ conclusion? The prior that is bestmatched to the data, in terms of likelihood, is the prior that sets p to f ≡ 140/250 with probability one. Let’s call this model H∗ . The likelihood ratio is P (D|H∗ )/P (D|H0 ) = 2250 f 140 (1 − f )110 = 6.1. So the strongest evidence that these data can possibly muster against the hypothesis that there is no bias is six-to-one.

While we are noticing the absurdly misleading answers that ‘sampling theory’ statistics produces, such as the p-value of 7% in the exercise we just solved, let’s stick the boot in. If we make a tiny change to the data set, increasing the number of heads in 250 tosses from 140 to 141, we find that the p-value goes below the mystical value of 0.05 (the p-value is 0.0497). The sampling theory statistician would happily squeak ‘the probability of getting a result as extreme as 141 heads is smaller than 0.05 – we thus reject the null hypothesis at a significance level of 5%’. The correct answer is shown for several values of α in figure 3.12. The values worth highlighting from this table are, first, the likelihood ratio when H1 uses the standard uniform prior, which is 1:0.61 in favour of the null hypothesis H0 . Second, the most favourable choice of α, from the point of view of H1 , can only yield a likelihood ratio of about 2.3:1 in favour of H1 . Be warned! A p-value of 0.05 is often interpreted as implying that the odds are stacked about twenty-to-one against the null hypothesis. But the truth in this case is that the evidence either slightly favours the null hypothesis, or disfavours it by at most 2.3 to one, depending on the choice of prior. The p-values and ‘significance levels’ of classical statistics should be treated with extreme caution. Shun them! Here ends the sermon.

α

P (D|H1 , α) P (D|H0 )

.37 1.0 2.7 7.4 20 55 148 403 1096

.25 .48 .82 1.3 1.8 1.9 1.7 1.3 1.1

Figure 3.11. Likelihood ratio for various choices of the prior distribution’s hyperparameter α.

α

P (D0 |H1 , α) P (D0 |H0 )

.37 1.0 2.7 7.4 20 55 148 403 1096

.32 .61 1.0 1.6 2.2 2.3 1.9 1.4 1.2

Figure 3.12. Likelihood ratio for various choices of the prior distribution’s hyperparameter α, when the data are D 0 = 141 heads in 250 trials.

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Part I

Data Compression

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About Chapter 4 In this chapter we discuss how to measure the information content of the outcome of a random experiment. This chapter has some tough bits. If you find the mathematical details hard, skim through them and keep going – you’ll be able to enjoy Chapters 5 and 6 without this chapter’s tools. Before reading Chapter 4, you should have read Chapter 2 and worked on exercises 2.21–2.25 and 2.16 (pp.36–37), and exercise 4.1 below. The following exercise is intended to help you think about how to measure information content. Exercise 4.1.[2, p.69] – Please work on this problem before reading Chapter 4. You are given 12 balls, all equal in weight except for one that is either heavier or lighter. You are also given a two-pan balance to use. In each use of the balance you may put any number of the 12 balls on the left pan, and the same number on the right pan, and push a button to initiate the weighing; there are three possible outcomes: either the weights are equal, or the balls on the left are heavier, or the balls on the left are lighter. Your task is to design a strategy to determine which is the odd ball and whether it is heavier or lighter than the others in as few uses of the balance as possible. While thinking about this problem, you may find it helpful to consider the following questions: (a) How can one measure information? (b) When you have identified the odd ball and whether it is heavy or light, how much information have you gained? (c) Once you have designed a strategy, draw a tree showing, for each of the possible outcomes of a weighing, what weighing you perform next. At each node in the tree, how much information have the outcomes so far given you, and how much information remains to be gained? (d) How much information is gained when you learn (i) the state of a flipped coin; (ii) the states of two flipped coins; (iii) the outcome when a four-sided die is rolled? (e) How much information is gained on the first step of the weighing problem if 6 balls are weighed against the other 6? How much is gained if 4 are weighed against 4 on the first step, leaving out 4 balls?

66

Notation x∈A S⊂A S⊆A V =B∪A V =B∩A |A|

x is a member of the set A S is a subset of the set A S is a subset of, or equal to, the set A V is the union of the sets B and A V is the intersection of the sets B and A number of elements in set A

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4 The Source Coding Theorem 4.1 How to measure the information content of a random variable? In the next few chapters, we’ll be talking about probability distributions and random variables. Most of the time we can get by with sloppy notation, but occasionally, we will need precise notation. Here is the notation that we established in Chapter 2. An ensemble X is a triple (x, AX , PX ), where the outcome x is the value of a random variable, which takes on one of a set of possible values, AX = {a1 , a2 , . . . , ai , . . . , aI }, having P probabilities PX = {p1 , p2 , . . . , pI }, with P (x = ai ) = pi , pi ≥ 0 and ai ∈AX P (x = ai ) = 1.

How can we measure the information content of an outcome x = a i from such an ensemble? In this chapter we examine the assertions 1. that the Shannon information content, 1 , (4.1) pi is a sensible measure of the information content of the outcome x = a i , and h(x = ai ) ≡ log2

2. that the entropy of the ensemble, H(X) =

X

pi log 2

i

1 , pi

(4.2)

is a sensible measure of the ensemble’s average information content. h(p) = log2

10 8

1 p

p

6 4 2 0 0

0.2

0.4

0.6

0.8

1

p

0.001 0.01 0.1 0.2 0.5

h(p) 10.0 6.6 3.3 2.3 1.0

H2 (p) 0.011 0.081 0.47 0.72 1.0

1

Figure 4.1. The Shannon information content h(p) = log2 1p and the binary entropy function H2 (p) = H(p, 1−p) = 1 p log2 p1 + (1 − p) log2 (1−p) as a function of p.

H2 (p)

0.8 0.6 0.4 0.2 0 0

0.2

0.4

0.6

0.8

1

p

Figure 4.1 shows the Shannon information content of an outcome with probability p, as a function of p. The less probable an outcome is, the greater its Shannon information content. Figure 4.1 also shows the binary entropy function, 1 1 H2 (p) = H(p, 1−p) = p log 2 + (1 − p) log 2 , (4.3) p (1 − p) which is the entropy of the ensemble X whose alphabet and probability distribution are AX = {a, b}, PX = {p, (1 − p)}. 67

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68

4 — The Source Coding Theorem

Information content of independent random variables Why should log 1/pi have anything to do with the information content? Why not some other function of pi ? We’ll explore this question in detail shortly, but first, notice a nice property of this particular function h(x) = log 1/p(x). Imagine learning the value of two independent random variables, x and y. The definition of independence is that the probability distribution is separable into a product: P (x, y) = P (x)P (y). (4.4) Intuitively, we might want any measure of the ‘amount of information gained’ to have the property of additivity – that is, for independent random variables x and y, the information gained when we learn x and y should equal the sum of the information gained if x alone were learned and the information gained if y alone were learned. The Shannon information content of the outcome x, y is h(x, y) = log

1 1 1 1 = log = log + log P (x, y) P (x)P (y) P (x) P (y)

(4.5)

so it does indeed satisfy h(x, y) = h(x) + h(y), if x and y are independent.

(4.6)

Exercise 4.2.[1, p.86] Show that, if x and y are independent, the entropy of the outcome x, y satisfies H(X, Y ) = H(X) + H(Y ).

(4.7)

In words, entropy is additive for independent variables. We now explore these ideas with some examples; then, in section 4.4 and in Chapters 5 and 6, we prove that the Shannon information content and the entropy are related to the number of bits needed to describe the outcome of an experiment.

The weighing problem: designing informative experiments Have you solved the weighing problem (exercise 4.1, p.66) yet? Are you sure? Notice that in three uses of the balance – which reads either ‘left heavier’, ‘right heavier’, or ‘balanced’ – the number of conceivable outcomes is 3 3 = 27, whereas the number of possible states of the world is 24: the odd ball could be any of twelve balls, and it could be heavy or light. So in principle, the problem might be solvable in three weighings – but not in two, since 3 2 < 24. If you know how you can determine the odd weight and whether it is heavy or light in three weighings, then you may read on. If you haven’t found a strategy that always gets there in three weighings, I encourage you to think about exercise 4.1 some more. Why is your strategy optimal? What is it about your series of weighings that allows useful information to be gained as quickly as possible? The answer is that at each step of an optimal procedure, the three outcomes (‘left heavier’, ‘right heavier’, and ‘balance’) are as close as possible to equiprobable. An optimal solution is shown in figure 4.2. Suboptimal strategies, such as weighing balls 1–6 against 7–12 on the first step, do not achieve all outcomes with equal probability: these two sets of balls can never balance, so the only possible outcomes are ‘left heavy’ and ‘right heavy’. Such a binary outcome rules out only half of the possible hypotheses,

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4.1: How to measure the information content of a random variable?

1+ 2+ 3+ 4+ 5+ 6+ 7+ 8+ 9+ 10+ 11+ 12+ 1− 2− 3− 4− 5− 6− 7− 8− 9− 10− 11− 12−

weigh 1234 5678

 B

  B B



 



 



 





-

B

B B

B

B B

B

B B

B

BNB

1+ 2+ 3+ 4+ 5− 6− 7− 8−

1− 2− 3− 4− 5+ 6+ 7+ 8+

9+ 10+ 11+ 12+ 9− 10− 11− 12−

weigh 126 345

weigh 126 345

weigh 9 10 11 123

   

  A A A A U A

   

  A A A A U A

   

  A A A A U A

1+ 2+ 5−

+ + −

3 4 6

− −

7 8

6+ 3− 4−

1− 2− 5+

+ +

7 8

69

1 2

 @ R @

3 4

 @ R @

1 7

 @ R @

3 4

 @ R @

1 2

 @ R @

7 1

 @ R @

1+ 2+ 5− 3+ 4+ 6− 7− 8− ? 4− 3− 6+ 2− 1− 5+ 7+ 8+ ? 9+

9+ 10+ 11+

9 10

 - 10+ @ R @ 11+

9− 10− 11−

9 10

10−  - 9− @ R @ 11−

12+  - 12− 12 12 @ R @ ? Figure 4.2. An optimal solution to the weighing problem. At each step there are two boxes: the left box shows which hypotheses are still possible; the right box shows the balls involved in the next weighing. The 24 hypotheses are written 1+ , . . . , 12− , with, e.g., 1+ denoting that 1 is the odd ball and it is heavy. Weighings are written by listing the names of the balls on the two pans, separated by a line; for example, in the first weighing, balls 1, 2, 3, and 4 are put on the left-hand side and 5, 6, 7, and 8 on the right. In each triplet of arrows the upper arrow leads to the situation when the left side is heavier, the middle arrow to the situation when the right side is heavier, and the lower arrow to the situation when the outcome is balanced. The three points labelled ? correspond to impossible outcomes. +



12 1

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70

4 — The Source Coding Theorem

so a strategy that uses such outcomes must sometimes take longer to find the right answer. The insight that the outcomes should be as near as possible to equiprobable makes it easier to search for an optimal strategy. The first weighing must divide the 24 possible hypotheses into three groups of eight. Then the second weighing must be chosen so that there is a 3:3:2 split of the hypotheses. Thus we might conclude: the outcome of a random experiment is guaranteed to be most informative if the probability distribution over outcomes is uniform. This conclusion agrees with the property of the entropy that you proved when you solved exercise 2.25 (p.37): the entropy of an ensemble X is biggest if all the outcomes have equal probability p i = 1/|AX |.

Guessing games In the game of twenty questions, one player thinks of an object, and the other player attempts to guess what the object is by asking questions that have yes/no answers, for example, ‘is it alive?’, or ‘is it human?’ The aim is to identify the object with as few questions as possible. What is the best strategy for playing this game? For simplicity, imagine that we are playing the rather dull version of twenty questions called ‘sixty-three’. Example 4.3. The game ‘sixty-three’. What’s the smallest number of yes/no questions needed to identify an integer x between 0 and 63? Intuitively, the best questions successively divide the 64 possibilities into equal sized sets. Six questions suffice. One reasonable strategy asks the following questions: 1: 2: 3: 4: 5: 6:

is is is is is is

x ≥ 32? x mod 32 ≥ 16? x mod 16 ≥ 8? x mod 8 ≥ 4? x mod 4 ≥ 2? x mod 2 = 1?

[The notation x mod 32, pronounced ‘x modulo 32’, denotes the remainder when x is divided by 32; for example, 35 mod 32 = 3 and 32 mod 32 = 0.] The answers to these questions, if translated from {yes, no} to {1, 0}, give the binary expansion of x, for example 35 ⇒ 100011. 2

What are the Shannon information contents of the outcomes in this example? If we assume that all values of x are equally likely, then the answers to the questions are independent and each has Shannon information content log 2 (1/0.5) = 1 bit; the total Shannon information gained is always six bits. Furthermore, the number x that we learn from these questions is a six-bit binary number. Our questioning strategy defines a way of encoding the random variable x as a binary file. So far, the Shannon information content makes sense: it measures the length of a binary file that encodes x. However, we have not yet studied ensembles where the outcomes have unequal probabilities. Does the Shannon information content make sense there too?

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4.1: How to measure the information content of a random variable? A B C D E F G H

×j

×j

1 2 3 4 5 6 7 8

move # question outcome

×

×

×××× ××× ×××× ××× ×j ××× ×××× × ×××× ××××

71 ×××××× × ×××××× ×××××× ×××××× ××××× ×j ×××××× ×××××× ×××××

×××××× × ×××××× ×××××× ×××××× ×××××× ×××××× ××××× ×j S ×××××

1 G3 x=n

2 B1 x=n

32 E5 x=n

48 F3 x=n

P (x)

63 64

62 63

32 33

16 17

1 16

h(x)

0.0227

0.0230

0.0443

0.0874

4.0

Total info.

0.0227

0.0458

1.0

2.0

6.0 Figure 4.3. A game of submarine. The submarine is hit on the 49th attempt.

The game of submarine: how many bits can one bit convey? In the game of battleships, each player hides a fleet of ships in a sea represented by a square grid. On each turn, one player attempts to hit the other’s ships by firing at one square in the opponent’s sea. The response to a selected square such as ‘G3’ is either ‘miss’, ‘hit’, or ‘hit and destroyed’. In a boring version of battleships called submarine, each player hides just one submarine in one square of an eight-by-eight grid. Figure 4.3 shows a few pictures of this game in progress: the circle represents the square that is being fired at, and the ×s show squares in which the outcome was a miss, x = n; the submarine is hit (outcome x = y shown by the symbol s) on the 49th attempt. Each shot made by a player defines an ensemble. The two possible outcomes are {y, n}, corresponding to a hit and a miss, and their probabilities depend on the state of the board. At the beginning, P (y) = 1/64 and P (n) = 63/64. At the second shot, if the first shot missed, P (y) = 1/63 and P (n) = 62/63. At the third shot, if the first two shots missed, P (y) = 1/62 and P (n) = 61/62. The Shannon information gained from an outcome x is h(x) = log(1/P (x)). If we are lucky, and hit the submarine on the first shot, then h(x) = h(1) (y) = log 2 64 = 6 bits.

(4.8)

Now, it might seem a little strange that one binary outcome can convey six bits. But we have learnt the hiding place, which could have been any of 64 squares; so we have, by one lucky binary question, indeed learnt six bits. What if the first shot misses? The Shannon information that we gain from this outcome is 64 = 0.0227 bits. (4.9) h(x) = h(1) (n) = log2 63 Does this make sense? It is not so obvious. Let’s keep going. If our second shot also misses, the Shannon information content of the second outcome is h(2) (n) = log2

63 = 0.0230 bits. 62

(4.10)

If we miss thirty-two times (firing at a new square each time), the total Shannon information gained is 64 63 33 + log 2 + · · · + log 2 63 62 32 = 0.0227 + 0.0230 + · · · + 0.0430 = 1.0 bits.

49 H3 x=y

log 2

(4.11)

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72

4 — The Source Coding Theorem

Why this round number? Well, what have we learnt? We now know that the submarine is not in any of the 32 squares we fired at; learning that fact is just like playing a game of sixty-three (p.70), asking as our first question ‘is x one of the thirty-two numbers corresponding to these squares I fired at?’, and receiving the answer ‘no’. This answer rules out half of the hypotheses, so it gives us one bit. After 48 unsuccessful shots, the information gained is 2 bits: the unknown location has been narrowed down to one quarter of the original hypothesis space. What if we hit the submarine on the 49th shot, when there were 16 squares left? The Shannon information content of this outcome is h(49) (y) = log2 16 = 4.0 bits.

(4.12)

The total Shannon information content of all the outcomes is 64 63 17 16 + log 2 + · · · + log 2 + log 2 63 62 16 1 = 0.0227 + 0.0230 + · · · + 0.0874 + 4.0 = 6.0 bits.

log 2

(4.13)

So once we know where the submarine is, the total Shannon information content gained is 6 bits. This result holds regardless of when we hit the submarine. If we hit it when there are n squares left to choose from – n was 16 in equation (4.13) – then the total information gained is: 64 63 n+1 n + log 2 + · · · + log 2 + log2 63 62 n   1 n+1 n 64 64 63 = log 2 × × ··· × × = log2 = 6 bits. (4.14) 63 62 n 1 1

log2

What have we learned from the examples so far? I think the submarine example makes quite a convincing case for the claim that the Shannon information content is a sensible measure of information content. And the game of sixty-three shows that the Shannon information content can be intimately connected to the size of a file that encodes the outcomes of a random experiment, thus suggesting a possible connection to data compression. In case you’re not convinced, let’s look at one more example.

The Wenglish language Wenglish is a language similar to English. Wenglish sentences consist of words drawn at random from the Wenglish dictionary, which contains 2 15 = 32,768 words, all of length 5 characters. Each word in the Wenglish dictionary was constructed at random by picking five letters from the probability distribution over a. . .z depicted in figure 2.1. Some entries from the dictionary are shown in alphabetical order in figure 4.4. Notice that the number of words in the dictionary (32,768) is much smaller than the total number of possible words of length 5 letters, 265 ' 12,000,000. Because the probability of the letter z is about 1/1000, only 32 of the words in the dictionary begin with the letter z. In contrast, the probability of the letter a is about 0.0625, and 2048 of the words begin with the letter a. Of those 2048 words, two start az, and 128 start aa. Let’s imagine that we are reading a Wenglish document, and let’s discuss the Shannon information content of the characters as we acquire them. If we

1 2 3

aaail aaaiu aaald .. .

129

abati .. .

2047 2048

azpan aztdn .. . .. . odrcr .. . .. . zatnt .. .

16 384

32 737 32 768

zxast

Figure 4.4. The Wenglish dictionary.

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4.2: Data compression

73

are given the text one word at a time, the Shannon information content of each five-character word is log 32,768 = 15 bits, since Wenglish uses all its words with equal probability. The average information content per character is therefore 3 bits. Now let’s look at the information content if we read the document one character at a time. If, say, the first letter of a word is a, the Shannon information content is log 1/0.0625 ' 4 bits. If the first letter is z, the Shannon information content is log 1/0.001 ' 10 bits. The information content is thus highly variable at the first character. The total information content of the 5 characters in a word, however, is exactly 15 bits; so the letters that follow an initial z have lower average information content per character than the letters that follow an initial a. A rare initial letter such as z indeed conveys more information about what the word is than a common initial letter. Similarly, in English, if rare characters occur at the start of the word (e.g. xyl...), then often we can identify the whole word immediately; whereas words that start with common characters (e.g. pro...) require more characters before we can identify them.

4.2 Data compression The preceding examples justify the idea that the Shannon information content of an outcome is a natural measure of its information content. Improbable outcomes do convey more information than probable outcomes. We now discuss the information content of a source by considering how many bits are needed to describe the outcome of an experiment. If we can show that we can compress data from a particular source into a file of L bits per source symbol and recover the data reliably, then we will say that the average information content of that source is at most L bits per symbol.

Example: compression of text files A file is composed of a sequence of bytes. A byte is composed of 8 bits and can have a decimal value between 0 and 255. A typical text file is composed of the ASCII character set (decimal values 0 to 127). This character set uses only seven of the eight bits in a byte. . Exercise 4.4.[1, p.86] By how much could the size of a file be reduced given that it is an ASCII file? How would you achieve this reduction? Intuitively, it seems reasonable to assert that an ASCII file contains 7/8 as much information as an arbitrary file of the same size, since we already know one out of every eight bits before we even look at the file. This is a simple example of redundancy. Most sources of data have further redundancy: English text files use the ASCII characters with non-equal frequency; certain pairs of letters are more probable than others; and entire words can be predicted given the context and a semantic understanding of the text.

Some simple data compression methods that define measures of information content One way of measuring the information content of a random variable is simply to count the number of possible outcomes, |A X |. (The number of elements in a set A is denoted by |A|.) If we gave a binary name to each outcome, the

Here we use the word ‘bit’ with its meaning, ‘a symbol with two values’, not to be confused with the unit of information content.

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74

4 — The Source Coding Theorem

length of each name would be log 2 |AX | bits, if |AX | happened to be a power of 2. We thus make the following definition. The raw bit content of X is H0 (X) = log2 |AX |.

(4.15)

H0 (X) is a lower bound for the number of binary questions that are always guaranteed to identify an outcome from the ensemble X. It is an additive quantity: the raw bit content of an ordered pair x, y, having |A X ||AY | possible outcomes, satisfies H0 (X, Y ) = H0 (X) + H0 (Y ). (4.16) This measure of information content does not include any probabilistic element, and the encoding rule it corresponds to does not ‘compress’ the source data, it simply maps each outcome to a constant-length binary string. Exercise 4.5.[2, p.86] Could there be a compressor that maps an outcome x to a binary code c(x), and a decompressor that maps c back to x, such that every possible outcome is compressed into a binary code of length shorter than H0 (X) bits? Even though a simple counting argument shows that it is impossible to make a reversible compression program that reduces the size of all files, amateur compression enthusiasts frequently announce that they have invented a program that can do this – indeed that they can further compress compressed files by putting them through their compressor several times. Stranger yet, patents have been granted to these modern-day alchemists. See the comp.compression frequently asked questions for further reading. 1 There are only two ways in which a ‘compressor’ can actually compress files: 1. A lossy compressor compresses some files, but maps some files to the same encoding. We’ll assume that the user requires perfect recovery of the source file, so the occurrence of one of these confusable files leads to a failure (though in applications such as image compression, lossy compression is viewed as satisfactory). We’ll denote by δ the probability that the source string is one of the confusable files, so a lossy compressor has a probability δ of failure. If δ can be made very small then a lossy compressor may be practically useful. 2. A lossless compressor maps all files to different encodings; if it shortens some files, it necessarily makes others longer. We try to design the compressor so that the probability that a file is lengthened is very small, and the probability that it is shortened is large. In this chapter we discuss a simple lossy compressor. In subsequent chapters we discuss lossless compression methods.

4.3 Information content defined in terms of lossy compression Whichever type of compressor we construct, we need somehow to take into account the probabilities of the different outcomes. Imagine comparing the information contents of two text files – one in which all 128 ASCII characters 1

http://sunsite.org.uk/public/usenet/news-faqs/comp.compression/

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4.3: Information content defined in terms of lossy compression

75

are used with equal probability, and one in which the characters are used with their frequencies in English text. Can we define a measure of information content that distinguishes between these two files? Intuitively, the latter file contains less information per character because it is more predictable. One simple way to use our knowledge that some symbols have a smaller probability is to imagine recoding the observations into a smaller alphabet – thus losing the ability to encode some of the more improbable symbols – and then measuring the raw bit content of the new alphabet. For example, we might take a risk when compressing English text, guessing that the most infrequent characters won’t occur, and make a reduced ASCII code that omits the characters { !, @, #, %, ^, *, ~, <, >, /, \, _, {, }, [, ], | }, thereby reducing the size of the alphabet by seventeen. The larger the risk we are willing to take, the smaller our final alphabet becomes. We introduce a parameter δ that describes the risk we are taking when using this compression method: δ is the probability that there will be no name for an outcome x. Example 4.6. Let AX = { a, b, c, d, e, f, g, h }, (4.17) 3 1 1 1 1 and PX = { 14 , 14 , 14 , 16 , 64 , 64 , 64 , 64 }. The raw bit content of this ensemble is 3 bits, corresponding to 8 binary names. But notice that P (x ∈ {a, b, c, d}) = 15/16. So if we are willing to run a risk of δ = 1/16 of not having a name for x, then we can get by with four names – half as many names as are needed if every x ∈ A X has a name. Table 4.5 shows binary names that could be given to the different outcomes in the cases δ = 0 and δ = 1/16. When δ = 0 we need 3 bits to encode the outcome; when δ = 1/16 we need only 2 bits. Let us now formalize this idea. To make a compression strategy with risk δ, we make the smallest possible subset S δ such that the probability that x is not in Sδ is less than or equal to δ, i.e., P (x 6∈ S δ ) ≤ δ. For each value of δ we can then define a new measure of information content – the log of the size of this smallest subset Sδ . [In ensembles in which several elements have the same probability, there may be several smallest subsets that contain different elements, but all that matters is their sizes (which are equal), so we will not dwell on this ambiguity.] The smallest δ-sufficient subset Sδ is the smallest subset of AX satisfying P (x ∈ Sδ ) ≥ 1 − δ.

(4.18)

The subset Sδ can be constructed by ranking the elements of A X in order of decreasing probability and adding successive elements starting from the most probable elements until the total probability is ≥ (1−δ). We can make a data compression code by assigning a binary name to each element of the smallest sufficient subset. This compression scheme motivates the following measure of information content: The essential bit content of X is: Hδ (X) = log 2 |Sδ |.

(4.19)

Note that H0 (X) is the special case of Hδ (X) with δ = 0 (if P (x) > 0 for all x ∈ AX ). [Caution: do not confuse H0 (X) and Hδ (X) with the function H2 (p) displayed in figure 4.1.] Figure 4.6 shows Hδ (X) for the ensemble of example 4.6 as a function of δ.

δ=0

δ = 1/16

x

c(x)

x

c(x)

a b c d e f g h

000 001 010 011 100 101 110 111

a b c d e f g h

00 01 10 11 − − − −

Table 4.5. Binary names for the outcomes, for two failure probabilities δ.

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76

4 — The Source Coding Theorem −6

−4

−2.4

S

S0

−2

log 2 P (x)

-

1 16

6

6

e,f,g,h

d

6 a,b,c

(a) {a,b,c,d,e,f,g,h} {a,b,c,d,e,f,g} {a,b,c,d,e,f} {a,b,c,d,e}

3 2.5

Hδ (X)

{a,b,c,d}

2

{a,b,c}

1.5 1

{a,b}

0.5 {a} 0 0

(b)

0.1

0.2

0.3

0.4

0.5

0.6

0.7

0.8

0.9

δ

Extended ensembles Is this compression method any more useful if we compress blocks of symbols from a source? We now turn to examples where the outcome x = (x 1 , x2 , . . . , xN ) is a string of N independent identically distributed random variables from a single ensemble X. We will denote by X N the ensemble (X1 , X2 , . . . , XN ). Remember that entropy is additive for independent variables (exercise 4.2 (p.68)), so H(X N ) = N H(X). Example 4.7. Consider a string of N flips of a bent coin, x = (x 1 , x2 , . . . , xN ), where xn ∈ {0, 1}, with probabilities p0 = 0.9, p1 = 0.1. The most probable strings x are those with most 0s. If r(x) is the number of 1s in x then N −r(x) r(x) P (x) = p0 p1 . (4.20) To evaluate Hδ (X N ) we must find the smallest sufficient subset S δ . This subset will contain all x with r(x) = 0, 1, 2, . . . , up to some r max (δ) − 1, and some of the x with r(x) = rmax (δ). Figures 4.7 and 4.8 show graphs of Hδ (X N ) against δ for the cases N = 4 and N = 10. The steps are the values of δ at which |Sδ | changes by 1, and the cusps where the slope of the staircase changes are the points where r max changes by 1. Exercise 4.8.[2, p.86] What are the mathematical shapes of the curves between the cusps? For the examples shown in figures 4.6–4.8, H δ (X N ) depends strongly on the value of δ, so it might not seem a fundamental or useful definition of information content. But we will consider what happens as N , the number of independent variables in X N , increases. We will find the remarkable result that Hδ (X N ) becomes almost independent of δ – and for all δ it is very close to N H(X), where H(X) is the entropy of one of the random variables. Figure 4.9 illustrates this asymptotic tendency for the binary ensemble of example 4.7. As N increases, N1 Hδ (X N ) becomes an increasingly flat function,

Figure 4.6. (a) The outcomes of X (from example 4.6 (p.75)), ranked by their probability. (b) The essential bit content Hδ (X). The labels on the graph show the smallest sufficient set as a function of δ. Note H0 (X) = 3 bits and H1/16 (X) = 2 bits.

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4.3: Information content defined in terms of lossy compression

−14

−12

−10

−8

−6

S0.01

6 1111

−4

log2 P (x) 0 -

S0.1

6

6

1101, 1011, . . .

−2

77

6

0110, 1010, . . .

6

0010, 0001, . . .

0000

Figure 4.7. (a) The sixteen outcomes of the ensemble X 4 with p1 = 0.1, ranked by probability. (b) The essential bit content Hδ (X 4 ). The upper schematic diagram indicates the strings’ probabilities by the vertical lines’ lengths (not to scale).

(a) 4 N=4 3.5

4

Hδ (X )

3 2.5 2 1.5 1 0.5 0

(b)

0

0.05

0.1

0.15

0.2

0.25

0.3

0.35

0.4

δ

10

Figure 4.8. Hδ (X N ) for N = 10 binary variables with p1 = 0.1.

N=10

Hδ (X 10 )

8

6

4

2

0 0

0.2

0.4

0.6

0.8

1

δ

1

1 N N Hδ (X )

Figure 4.9. N1 Hδ (X N ) for N = 10, 210, . . . , 1010 binary variables with p1 = 0.1.

N=10 N=210 N=410 N=610 N=810 N=1010

0.8

0.6

0.4

0.2

0 0

0.2

0.4

0.6

0.8

1

δ

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78

4 — The Source Coding Theorem

x

log 2 (P (x))

...1...................1.....1....1.1.......1........1...........1.....................1.......11... ......................1.....1.....1.......1....1.........1.....................................1.... ........1....1..1...1....11..1.1.........11.........................1...1.1..1...1................1. 1.1...1................1.......................11.1..1............................1.....1..1.11..... ...11...........1...1.....1.1......1..........1....1...1.....1............1......................... ..............1......1.........1.1.......1..........1............1...1......................1....... .....1........1.......1...1............1............1...........1......1..11........................ .....1..1..1...............111...................1...............1.........1.1...1...1.............1 .........1..........1.....1......1..........1....1..............................................1... ......1........................1..............1.....1..1.1.1..1...................................1. 1.......................1..........1...1...................1....1....1........1..11..1.1...1........ ...........11.1.........1................1......1.....................1............................. .1..........1...1.1.............1.......11...........1.1...1..............1.............11.......... ......1...1..1.....1..11.1.1.1...1.....................1............1.............1..1.............. ............11.1......1....1..1............................1.......1..............1.......1.........

.................................................................................................... 1111111111111111111111111111111111111111111111111111111111111111111111111111111111111111111111111111

−50.1 −37.3 −65.9 −56.4 −53.2 −43.7 −46.8 −56.4 −37.3 −43.7 −56.4 −37.3 −56.4 −59.5 −46.8 −15.2 −332.1

except for tails close to δ = 0 and 1. As long as we are allowed a tiny probability of error δ, compression down to N H bits is possible. Even if we are allowed a large probability of error, we still can compress only down to N H bits. This is the source coding theorem. Theorem 4.1 Shannon’s source coding theorem. Let X be an ensemble with entropy H(X) = H bits. Given  > 0 and 0 < δ < 1, there exists a positive integer N0 such that for N > N0 , 1 Hδ (X N ) − H < . (4.21) N

4.4 Typicality

Why does increasing N help? Let’s examine long strings from X N . Table 4.10 shows fifteen samples from X N for N = 100 and p1 = 0.1. The probability of a string x that contains r 1s and N −r 0s is P (x) = pr1 (1 − p1 )N −r . The number of strings that contain r 1s is   N n(r) = . r So the number of 1s, r, has a binomial distribution:   N r p (1 − p1 )N −r . P (r) = r 1

(4.22)

(4.23)

(4.24)

These functions are shown in figure 4.11. The mean of r is N p 1 , and its p standard deviation is N p1 (1 − p1 ) (p.1). If N is 100 then r ∼ N p1 ±

p

N p1 (1 − p1 ) ' 10 ± 3.

(4.25)

Figure 4.10. The top 15 strings are samples from X 100 , where p1 = 0.1 and p0 = 0.9. The bottom two are the most and least probable strings in this ensemble. The final column shows the log-probabilities of the random strings, which may be compared with the entropy H(X 100 ) = 46.9 bits.

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4.4: Typicality

79

N = 100

n(r) =

N r



N = 1000

1.2e+29

3e+299

1e+29

2.5e+299

8e+28

2e+299

6e+28

1.5e+299

4e+28

1e+299

2e+28

5e+298

0

0 0

10 20 30 40 50 60 70 80 90 100

0 100 200 300 400 500 600 700 800 9001000

2e-05

P (x) = pr1 (1 − p1 )N −r

2e-05 1e-05

0

1e-05

0

1

2

3

4

5

0 0

log2 P (x)

10 20 30 40 50 60 70 80 90 100

0

0

-50

-500

-100

-1000

T

-150

-1500

-200

-2000

-250

-2500

-300

-3000

-350

-3500 0

10 20 30 40 50 60 70 80 90 100

0.14

N r



pr1 (1 − p1 )N −r

0 100 200 300 400 500 600 700 800 9001000 0.045 0.04 0.035 0.03 0.025 0.02 0.015 0.01 0.005 0

0.12

n(r)P (x) =

T

0.1 0.08 0.06 0.04 0.02 0 0

10 20 30 40 50 60 70 80 90 100

r

0 100 200 300 400 500 600 700 800 9001000

r

Figure 4.11. Anatomy of the typical set T . For p1 = 0.1 and N = 100 and N = 1000, these graphs show n(r), the number of strings containing r 1s; the probability P (x) of a single string that contains r 1s; the same probability on a log scale; and the total probability n(r)P (x) of all strings that contain r 1s. The number r is on the horizontal axis. The plot of log 2 P (x) also shows by a dotted line the mean value of log2 P (x) = −N H2 (p1 ), which equals −46.9 when N = 100 and −469 when N = 1000. The typical set includes only the strings that have log2 P (x) close to this value. The range marked T shows the set TN β (as defined in section 4.4) for N = 100 and β = 0.29 (left) and N = 1000, β = 0.09 (right).

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80

4 — The Source Coding Theorem

If N = 1000 then r ∼ 100 ± 10.

(4.26)

Notice that as N gets bigger, the probability distribution of r becomes more concentrated, in the sense that while the range √ of possible values of r grows as N , the standard deviation of r grows only as N . That r is most likely to fall in a small range of values implies that the outcome x is also most likely to fall in a corresponding small subset of outcomes that we will call the typical set.

Definition of the typical set Let us define typicality for an arbitrary ensemble X with alphabet A X . Our definition of a typical string will involve the string’s probability. A long string of N symbols will usually contain about p 1 N occurrences of the first symbol, p2 N occurrences of the second, etc. Hence the probability of this string is roughly (p1 N ) (p2 N ) p2

P (x)typ = P (x1 )P (x2 )P (x3 ) . . . P (xN ) ' p1

so that the information content of a typical string is X 1 1 log2 ' N = N H. pi log2 P (x) pi

(pI N )

. . . pI

(4.27)

(4.28)

i

So the random variable log 2 1/P (x), which is the information content of x, is very likely to be close in value to N H. We build our definition of typicality on this observation. We define the typical elements of AN X to be those elements that have prob−N H ability close to 2 . (Note that the typical set, unlike the smallest sufficient subset, does not include the most probable elements of A N X , but we will show that these most probable elements contribute negligible probability.) We introduce a parameter β that defines how close the probability has to be to 2−N H for an element to be ‘typical’. We call the set of typical elements the typical set, TN β :   1 N 1 − H < β . (4.29) TN β ≡ x ∈ AX : log2 N P (x) We will show that whatever value of β we choose, the typical set contains almost all the probability as N increases. This important result is sometimes called the ‘asymptotic equipartition’ principle.

‘Asymptotic equipartition’ principle. For an ensemble of N independent identically distributed (i.i.d.) random variables X N ≡ (X1 , X2 , . . . , XN ), with N sufficiently large, the outcome x = (x 1 , x2 , . . . , xN ) is almost N H(X) members, each certain to belong to a subset of AN X having only 2 −N H(X) having probability ‘close to’ 2 . Notice that if H(X) < H0 (X) then 2N H(X) is a tiny fraction of the number N N H0 (X) . of possible outcomes |AN X | = |AX | = 2 The term equipartition is chosen to describe the idea that the members of the typical set have roughly equal probability. [This should not be taken too literally, hence my use of quotes around ‘asymptotic equipartition’; see page 83.] A second meaning for equipartition, in thermal physics, is the idea that each degree of freedom of a classical system has equal average energy, 12 kT . This second meaning is not intended here.

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4.5: Proofs

81 log2 P (x) −N H(X)

-

TN β

6

6

6

6

6

1111111111110. . . 11111110111 0000100000010. . . 00001000010 0100000001000. . . 00010000000 0001000000000. . . 00000000000 0000000000000. . . 00000000000

The ‘asymptotic equipartition’ principle is equivalent to: Shannon’s source coding theorem (verbal statement). N i.i.d. random variables each with entropy H(X) can be compressed into more than N H(X) bits with negligible risk of information loss, as N → ∞; conversely if they are compressed into fewer than N H(X) bits it is virtually certain that information will be lost. These two theorems are equivalent because we can define a compression algorithm that gives a distinct name of length N H(X) bits to each x in the typical set.

4.5 Proofs This section may be skipped if found tough going.

The law of large numbers Our proof of the source coding theorem uses the law of large numbers. P Mean and variance of a real random variable are E[u] = u ¯ = u P (u)u P and var(u) = σu2 = E[(u − u ¯)2 ] = u P (u)(u − u ¯)2 .

Technical note: strictly I am assuming here that u is a function u(x) of P a sample x from a finite discrete P ensemble X. Then the summations u P (u)f (u) should be written x P (x)f (u(x)). This means that P (u) is a finite sum of delta functions. This restriction guarantees that the mean and variance of u do exist, which is not necessarily the case for general P (u).

Chebyshev’s inequality 1. Let t be a non-negative real random variable, and let α be a positive real number. Then P (t ≥ α) ≤

t¯ . α

(4.30)

P Proof: P (t ≥ α) = P t≥α P (t). We multiply each term by t/α ≥ 1 and obtain: P (t ≥ α) ≤ t≥α P (t)t/α. P We add the (non-negative) missing 2 terms and obtain: P (t ≥ α) ≤ t P (t)t/α = t¯/α.

Figure 4.12. Schematic diagram showing all strings in the ensemble X N ranked by their probability, and the typical set TN β .

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82

4 — The Source Coding Theorem

Chebyshev’s inequality 2. Let x be a random variable, and let α be a positive real number. Then  P (x − x ¯)2 ≥ α ≤ σx2 /α. (4.31) Proof: Take t = (x − x ¯)2 and apply the previous proposition.

2

Weak law of large numbers. Take x to be the average of N independent ¯ and common varirandom variablesPh1 , . . . , hN , having common mean h h . Then ance σh2 : x = N1 N n=1 n ¯ 2 ≥ α) ≤ σ 2 /αN. P ((x − h) (4.32) h

¯ and that σ 2 = σ 2 /N . Proof: obtained by showing that x ¯=h x h

2

We are interested in x being very close to the mean (α very small). No matter how large σh2 is, and no matter how small the required α is, and no matter ¯ 2 ≥ α, we can always achieve it how small the desired probability that (x − h) by taking N large enough.

Proof of theorem 4.1 (p.78) 1 We apply the law of large numbers to the random variable N1 log2 P (x) defined N for x drawn from the ensemble X . This random variable can be written as the average of N information contents h n = log 2 (1/P (xn )), each of which is a random variable with mean H = H(X) and variance σ 2 ≡ var[log 2 (1/P (xn ))]. (Each term hn is the Shannon information content of the nth outcome.) We again define the typical set with parameters N and β thus: ) ( 2  1 1 (4.33) log 2 − H < β2 . TN β = x ∈ A N X : N P (x)

For all x ∈ TN β , the probability of x satisfies

2−N (H+β) < P (x) < 2−N (H−β) .

(4.34)

And by the law of large numbers, σ2 . (4.35) β2N We have thus proved the ‘asymptotic equipartition’ principle. As N increases, the probability that x falls in TN β approaches 1, for any β. How does this result relate to source coding? We must relate TN β to Hδ (X N ). We will show that for any given δ there is a sufficiently big N such that Hδ (X N ) ' N H. P (x ∈ TN β ) ≥ 1 −

Part 1:

1 N N Hδ (X )

1 N

Hδ (X N )

H0 (X)

< H + .

The set TN β is not the best subset for compression. So the size of T N β gives an upper bound on Hδ . We show how small Hδ (X N ) must be by calculating how big TN β could possibly be. We are free to set β to any convenient value. The smallest possible probability that a member of T N β can have is 2−N (H+β) , and the total probability contained by T N β can’t be any bigger than 1. So |TN β | 2

−N (H+β)

< 1,

σ2 2 N

H H− 0

1

δ

(4.36)

that is, the size of the typical set is bounded by |TN β | < 2N (H+β) .

H+

(4.37)

≤ δ, then P (TN β ) ≥ 1 − δ, and the set If we set β =  and N0 such that 0 TN β becomes a witness to the fact that Hδ (X N ) ≤ log 2 |TN β | < N (H + ).

Figure 4.13. Schematic illustration of the two parts of the theorem. Given any δ and , we show that for large enough N , N1 Hδ (X N ) lies (1) below the line H +  and (2) above the line H − .

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4.6: Comments Part 2:

1 N N Hδ (X )

83 > H − .

Imagine that someone claims this second part is not so – that, for any N , the smallest δ-sufficient subset Sδ is smaller than the above inequality would allow. We can make use of our typical set to show that they must be mistaken. Remember that we are free to set β to any value we choose. We will set β = /2, so that our task is to prove that a subset S 0 having |S 0 | ≤ 2N (H−2β) and achieving P (x ∈ S 0 ) ≥ 1 − δ cannot exist (for N greater than an N 0 that we will specify). So, let us consider the probability of falling in this rival smaller subset S 0 . The probability of the subset S 0 is 0

0

0

P (x ∈ S ) = P (x ∈ S ∩TN β ) + P (x ∈ S ∩TN β ),

(4.38)

where TN β denotes the complement {x 6∈ TN β }. The maximum value of the first term is found if S 0 ∩ TN β contains 2N (H−2β) outcomes all with the maximum probability, 2−N (H−β) . The maximum value the second term can have is P (x 6∈ TN β ). So: P (x ∈ S 0 ) ≤ 2N (H−2β) 2−N (H−β) +

σ2 σ2 = 2−N β + 2 . 2 β N β N

(4.39)

We can now set β = /2 and N0 such that P (x ∈ S 0 ) < 1 − δ, which shows that S 0 cannot satisfy the definition of a sufficient subset S δ . Thus any subset S 0 with size |S 0 | ≤ 2N (H−) has probability less than 1 − δ, so by the definition of Hδ , Hδ (X N ) > N (H − ). Thus for large enough N , the function N1 Hδ (X N ) is essentially a constant function of δ, for 0 < δ < 1, as illustrated in figures 4.9 and 4.13. 2

4.6 Comments The source coding theorem (p.78) has two parts, N1 Hδ (X N ) < H + , and 1 N N Hδ (X ) > H − . Both results are interesting. The first part tells us that even if the probability of error δ is extremely small, the number of bits per symbol N1 Hδ (X N ) needed to specify a long N -symbol string x with vanishingly small error probability does not have to exceed H +  bits. We need to have only a tiny tolerance for error, and the number of bits required drops significantly from H 0 (X) to (H + ). What happens if we are yet more tolerant to compression errors? Part 2 tells us that even if δ is very close to 1, so that errors are made most of the time, the average number of bits per symbol needed to specify x must still be at least H −  bits. These two extremes tell us that regardless of our specific allowance for error, the number of bits per symbol needed to specify x is H bits; no more and no less.

Caveat regarding ‘asymptotic equipartition’ I put the words ‘asymptotic equipartition’ in quotes because it is important not to think that the elements of the typical set T N β really do have roughly the same probability as each other. They are similar in probability only in 1 the sense that their values of log 2 P (x) are within 2N β of each other. Now, as β is decreased, how does N have to increase, if we are to keep our bound on 2 the mass of the typical set, P (x ∈ TN β ) ≥ 1 − βσ2 N , constant? N must grow √ as 1/β 2 , so, if we write β in terms of N as α/ N , for some constant α, then

TN β

'$ '$ S0

I @ OCC @ S0 ∩ T &% &% Nβ C 0 S ∩ TN β

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84

4 — The Source Coding Theorem √

the most probable string in the typical set will be of order 2 α N times greater than the least probable string in the typical set. As β decreases, N increases, √ α N and this ratio 2 grows exponentially. Thus we have ‘equipartition’ only in a weak sense!

Why did we introduce the typical set? The best choice of subset for block compression is (by definition) S δ , not a typical set. So why did we bother introducing the typical set? The answer is, we can count the typical set. We know that all its elements have ‘almost identical’ probability (2−N H ), and we know the whole set has probability almost 1, so the typical set must have roughly 2 N H elements. Without the help of the typical set (which is very similar to S δ ) it would have been hard to count how many elements there are in Sδ .

4.7 Exercises Weighing problems . Exercise 4.9.[1 ] While some people, when they first encounter the weighing problem with 12 balls and the three-outcome balance (exercise 4.1 (p.66)), think that weighing six balls against six balls is a good first weighing, others say ‘no, weighing six against six conveys no information at all’. Explain to the second group why they are both right and wrong. Compute the information gained about which is the odd ball, and the information gained about which is the odd ball and whether it is heavy or light. . Exercise 4.10.[2 ] Solve the weighing problem for the case where there are 39 balls of which one is known to be odd. . Exercise 4.11.[2 ] You are given 16 balls, all of which are equal in weight except for one that is either heavier or lighter. You are also given a bizarre twopan balance that can report only two outcomes: ‘the two sides balance’ or ‘the two sides do not balance’. Design a strategy to determine which is the odd ball in as few uses of the balance as possible. . Exercise 4.12.[2 ] You have a two-pan balance; your job is to weigh out bags of flour with integer weights 1 to 40 pounds inclusive. How many weights do you need? [You are allowed to put weights on either pan. You’re only allowed to put one flour bag on the balance at a time.] Exercise 4.13.[4, p.86] (a) Is it possible to solve exercise 4.1 (p.66) (the weighing problem with 12 balls and the three-outcome balance) using a sequence of three fixed weighings, such that the balls chosen for the second weighing do not depend on the outcome of the first, and the third weighing does not depend on the first or second? (b) Find a solution to the general N -ball weighing problem in which exactly one of N balls is odd. Show that in W weighings, an odd ball can be identified from among N = (3 W − 3)/2 balls. Exercise 4.14.[3 ] You are given 12 balls and the three-outcome balance of exercise 4.1; this time, two of the balls are odd; each odd ball may be heavy or light, and we don’t know which. We want to identify the odd balls and in which direction they are odd.

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4.7: Exercises

85

(a) Estimate how many weighings are required by the optimal strategy. And what if there are three odd balls? (b) How do your answers change if it is known that all the regular balls weigh 100 g, that light balls weigh 99 g, and heavy ones weigh 110 g?

Source coding with a lossy compressor, with loss δ . Exercise 4.15.[2, p.87] Let PX = {0.2, 0.8}. Sketch δ for N = 1, 2 and 1000. . Exercise 4.16.[2 ] Let PY = {0.5, 0.5}. Sketch N = 1, 2, 3 and 100.

1 N N Hδ (X )

1 N N Hδ (Y )

as a function of

as a function of δ for

. Exercise 4.17.[2, p.87] (For physics students.) Discuss the relationship between the proof of the ‘asymptotic equipartition’ principle and the equivalence (for large systems) of the Boltzmann entropy and the Gibbs entropy.

Distributions that don’t obey the law of large numbers The law of large numbers, which we used in this chapter, shows that the mean of a set of N i.i.d. random variables has a probability distribution that becomes √ narrower, with width ∝ 1/ N , as N increases. However, we have proved this property only for discrete random variables, that is, for real numbers taking on a finite set of possible values. While many random variables with continuous probability distributions also satisfy the law of large numbers, there are important distributions that do not. Some continuous distributions do not have a mean or variance. . Exercise 4.18.[3, p.88] Sketch the Cauchy distribution P (x) =

1 1 , Z x2 + 1

x ∈ (−∞, ∞).

(4.40)

What is its normalizing constant Z? Can you evaluate its mean or variance? Consider the sum z = x1 + x2 , where x1 and x2 are independent random variables from a Cauchy distribution. What is P (z)? What is the probability distribution of the mean of x 1 and x2 , x ¯ = (x1 + x2 )/2? What is the probability distribution of the mean of N samples from this Cauchy distribution?

Other asymptotic properties Exercise 4.19.[3 ] Chernoff bound. We derived the weak law of large numbers from Chebyshev’s inequality (4.30) by letting the random variable t in the inequality P (t ≥ α) ≤ t¯/α be a function, t = (x− x ¯)2 , of the random variable x we were interested in. Other useful inequalities can be obtained by using other functions. The Chernoff bound, which is useful for bounding the tails of a distribution, is obtained by letting t = exp(sx). Show that P (x ≥ a) ≤ e−sa g(s),

for any s > 0

(4.41)

P (x ≤ a) ≤ e−sa g(s),

for any s < 0

(4.42)

and

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86

4 — The Source Coding Theorem where g(s) is the moment-generating function of x, X g(s) = P (x) esx .

(4.43)

x

Curious functions related to p log 1/p Exercise 4.20.[4, p.89] This exercise has no purpose at all; it’s included for the enjoyment of those who like mathematical curiosities. Sketch the function f (x) = x

xx

xx

··

·

(4.44)

for x ≥ 0. Hint: Work out the inverse function to f – that is, the function g(y) such that if x = g(y) then y = f (x) – it’s closely related to p log 1/p.

4.8 Solutions Solution to exercise 4.2 (p.68).

Let P (x, y) = P (x)P (y). Then

1 (4.45) P (x)P (y) xy X X 1 1 = P (x)P (y) log + (4.46) P (x)P (y) log P (x) P (y) xy xy X X 1 1 = + (4.47) P (x) log P (y) log P (x) P (y) x y

H(X, Y ) =

X

P (x)P (y) log

= H(X) + H(Y ).

(4.48)

Solution to exercise 4.4 (p.73). An ASCII file can be reduced in size by a factor of 7/8. This reduction could be achieved by a block code that maps 8-byte blocks into 7-byte blocks by copying the 56 information-carrying bits into 7 bytes, and ignoring the last bit of every character. Solution to exercise 4.5 (p.74). The pigeon-hole principle states: you can’t put 16 pigeons into 15 holes without using one of the holes twice. Similarly, you can’t give AX outcomes unique binary names of some length l shorter than log 2 |AX | bits, because there are only 2l such binary names, and l < log 2 |AX | implies 2l < |AX |, so at least two different inputs to the compressor would compress to the same output file. Solution to exercise 4.8 (p.76). Between the cusps, all the changes in probability are equal, and the number of elements in T changes by one at each step. So Hδ varies logarithmically with (−δ). Solution to exercise 4.13 (p.84). This solution was found by Dyson and Lyness in 1946 and presented in the following elegant form by John Conway in 1999. Be warned: the symbols A, B, and C are used to name the balls, to name the pans of the balance, to name the outcomes, and to name the possible states of the odd ball! (a) Label the 12 balls by the sequences AAB

ABA

and in the

ABB

ABC

BBC

BCA

BCB

BCC

CAA

CAB

CAC

CCA

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4.8: Solutions

87

1st AAB ABA ABB ABC BBC BCA BCB BCC 2nd weighings put AAB CAA CAB CAC in pan A, ABA ABB ABC BBC in pan B. 3rd ABA BCA CAA CCA AAB ABB BCB CAB

Now in a given weighing, a pan will either end up in the • Canonical position (C) that it assumes when the pans are balanced, or • Above that position (A), or

• Below it (B),

so the three weighings determine for each pan a sequence of three of these letters. If both sequences are CCC, then there’s no odd ball. Otherwise, for just one of the two pans, the sequence is among the 12 above, and names the odd ball, whose weight is Above or Below the proper one according as the pan is A or B. (b) In W weighings the odd ball can be identified from among N = (3W − 3)/2

(4.49)

balls in the same way, by labelling them with all the non-constant sequences of W letters from A, B, C whose first change is A-to-B or B-to-C or C-to-A, and at the wth weighing putting those whose wth letter is A in pan A and those whose wth letter is B in pan B. Solution to exercise 4.15 (p.85). The curves N1 Hδ (X N ) as a function of δ for N = 1, 2 and 1000 are shown in figure 4.14. Note that H 2 (0.2) = 0.72 bits. 1

N =1

N=1 N=2 N=1000 0.8

0.6

0.4

N =2

δ

1 N Hδ (X)

2Hδ (X)

δ

1 N Hδ (X)

2Hδ (X)

0–0.2 0.2–1

1 0

2 1

0–0.04 0.04–0.2 0.2–0.36 0.36–1

1 0.79 0.5 0

4 3 2 1

0.2

0 0

0.2

0.4

0.6

0.8

1

P Solution to exercise 4.17 (p.85). The Gibbs entropy is k B i pi ln p1i , where i runs over all states of the system. This entropy is equivalent (apart from the factor of kB ) to the Shannon entropy of the ensemble. Whereas the Gibbs entropy can be defined for any ensemble, the Boltzmann entropy is only defined for microcanonical ensembles, which have a probability distribution that is uniform over a set of accessible states. The Boltzmann entropy is defined to be SB = kB ln Ω where Ω is the number of accessible states of the microcanonical ensemble. This is equivalent (apart from the factor of kB ) to the perfect information content H 0 of that constrained ensemble. The Gibbs entropy of a microcanonical ensemble is trivially equal to the Boltzmann entropy.

Figure 4.14. N1 Hδ (X) (vertical axis) against δ (horizontal), for N = 1, 2, 100 binary variables with p1 = 0.4.

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88

4 — The Source Coding Theorem

We now consider a thermal distribution (the canonical ensemble), where the probability of a state x is   1 E(x) P (x) = exp − . (4.50) Z kB T With this canonical ensemble we can associate a corresponding microcanonical ensemble, an ensemble with total energy fixed to the mean energy of the canonical ensemble (fixed to within some precision ). Now, fixing the total energy to a precision  is equivalent to fixing the value of ln 1/P (x) to within kB T . Our definition of the typical set T N β was precisely that it consisted of all elements that have a value of log P (x) very close to the mean value of log P (x) under the canonical ensemble, −N H(X). Thus the microcanonical ensemble is equivalent to a uniform distribution over the typical set of the canonical ensemble. Our proof of the ‘asymptotic equipartition’ principle thus proves – for the case of a system whose energy is separable into a sum of independent terms – that the Boltzmann entropy of the microcanonical ensemble is very close (for large N ) to the Gibbs entropy of the canonical ensemble, if the energy of the microcanonical ensemble is constrained to equal the mean energy of the canonical ensemble. Solution to exercise 4.18 (p.85). The normalizing constant of the Cauchy distribution 1 1 P (x) = Z x2 + 1 is Z ∞  ∞ π −π 1 Z= = tan−1 x −∞ = − = π. (4.51) dx 2 x +1 2 2 −∞

The mean and variance of this distribution are both undefined. (The distribution is symmetrical about zero, but this does not imply that its mean is zero. The mean is the value of a divergent integral.) The sum z = x 1 + x2 , where x1 and x2 both have Cauchy distributions, has probability density given by the convolution Z ∞ 1 1 1 P (z) = 2 , (4.52) dx1 2 π −∞ x1 + 1 (z − x1 )2 + 1

which after a considerable labour using standard methods gives P (z) =

1 1 π 2 2 2 = , 2 2 π z +4 π z + 22

(4.53)

which we recognize as a Cauchy distribution with width parameter 2 (where the original distribution has width parameter 1). This implies that the mean of the two points, x ¯ = (x1 + x2 )/2 = z/2, has a Cauchy distribution with width parameter 1. Generalizing, the mean of N samples from a Cauchy distribution is Cauchy-distributed with the same parameters as the individual samples. √ The probability distribution of the mean does not become narrower as 1/ N . The central-limit theorem does not apply to the Cauchy distribution, because it does not have a finite variance. An alternative neat method for getting to equation (4.53) makes use of the Fourier transform of the Cauchy distribution, which is a biexponential e −|ω| . Convolution in real space corresponds to multiplication in Fourier space, so the Fourier transform of z is simply e −|2ω| . Reversing the transform, we obtain equation (4.53).

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4.8: Solutions

89

Solution to exercise 4.20 (p.86).

The function f (x) has inverse function 50

g(y) = y

1/y

.

(4.54)

40 30

Note

20

log g(y) = 1/y log y.

(4.55)

I obtained a tentative graph of f (x) by plotting g(y) with y along the vertical axis and g(y) along the horizontal axis. The resulting graph suggests that f (x) is single valued for x ∈ (0, 1), and looks surprisingly well-behaved and √ ordinary; for x ∈ (1, e1/e ), f (x) is two-valued. f ( 2) is equal both to 2 and 4. For x > e1/e (which is about 1.44), f (x) is infinite. However, it might be argued that this approach to sketching f (x) is only partly valid, if we define f x as the limit of the sequence of functions x, x x , xx , . . .; this sequence does not e have a limit for 0 ≤ x ≤ (1/e) ' 0.07 on account of a pitchfork bifurcation at x = (1/e)e ; and for x ∈ (1, e1/e ), the sequence’s limit is single-valued – the lower of the two values sketched in the figure.

10 0 0

0.2

0.4

0.6

0.8

1

1.2

1.4

0.8

1

1.2

1.4

5 4 3 2 1 0 0

0.2

0.4

0.6

0.5 0.4 0.3 0.2 0.1 0 0

0.2

x xx,

xx

Figure 4.15. f (x) = at three different scales.

··

·

shown

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About Chapter 5 In the last chapter, we saw a proof of the fundamental status of the entropy as a measure of average information content. We defined a data compression scheme using fixed length block codes, and proved that as N increases, it is possible to encode N i.i.d. variables x = (x 1 , . . . , xN ) into a block of N (H(X)+ ) bits with vanishing probability of error, whereas if we attempt to encode X N into N (H(X) − ) bits, the probability of error is virtually 1. We thus verified the possibility of data compression, but the block coding defined in the proof did not give a practical algorithm. In this chapter and the next, we study practical data compression algorithms. Whereas the last chapter’s compression scheme used large blocks of fixed size and was lossy, in the next chapter we discuss variable-length compression schemes that are practical for small block sizes and that are not lossy. Imagine a rubber glove filled with water. If we compress two fingers of the glove, some other part of the glove has to expand, because the total volume of water is constant. (Water is essentially incompressible.) Similarly, when we shorten the codewords for some outcomes, there must be other codewords that get longer, if the scheme is not lossy. In this chapter we will discover the information-theoretic equivalent of water volume. Before reading Chapter 5, you should have worked on exercise 2.26 (p.37). We will use the following notation for intervals: x ∈ [1, 2) x ∈ (1, 2]

means that x ≥ 1 and x < 2; means that x > 1 and x ≤ 2.

90

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5 Symbol Codes In this chapter, we discuss variable-length symbol codes, which encode one source symbol at a time, instead of encoding huge strings of N source symbols. These codes are lossless: unlike the last chapter’s block codes, they are guaranteed to compress and decompress without any errors; but there is a chance that the codes may sometimes produce encoded strings longer than the original source string. The idea is that we can achieve compression, on average, by assigning shorter encodings to the more probable outcomes and longer encodings to the less probable. The key issues are: What are the implications if a symbol code is lossless? If some codewords are shortened, by how much do other codewords have to be lengthened? Making compression practical. How can we ensure that a symbol code is easy to decode? Optimal symbol codes. How should we assign codelengths to achieve the best compression, and what is the best achievable compression? We again verify the fundamental status of the Shannon information content and the entropy, proving: Source coding theorem (symbol codes). There exists a variable-length encoding C of an ensemble X such that the average length of an encoded symbol, L(C, X), satisfies L(C, X) ∈ [H(X), H(X) + 1). The average length is equal to the entropy H(X) only if the codelength for each outcome is equal to its Shannon information content. We will also define a constructive procedure, the Huffman coding algorithm, that produces optimal symbol codes. Notation for alphabets. AN denotes the set of ordered N -tuples of elements from the set A, i.e., all strings of length N . The symbol A + will denote the set of all strings of finite length composed of elements from the set A. Example 5.1. {0, 1}3 = {000, 001, 010, 011, 100, 101, 110, 111}. Example 5.2. {0, 1}+ = {0, 1, 00, 01, 10, 11, 000, 001, . . .}. 91

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92

5 — Symbol Codes

5.1 Symbol codes A (binary) symbol code C for an ensemble X is a mapping from the range of x, AX = {a1 , . . . , aI }, to {0, 1}+ . c(x) will denote the codeword corresponding to x, and l(x) will denote its length, with l i = l(ai ). + The extended code C + is a mapping from A+ X to {0, 1} obtained by concatenation, without punctuation, of the corresponding codewords:

c+ (x1 x2 . . . xN ) = c(x1 )c(x2 ) . . . c(xN ).

(5.1)

[The term ‘mapping’ here is a synonym for ‘function’.] Example 5.3. A symbol code for the ensemble X defined by AX PX

= { a, b, c, d }, = { 1/2, 1/4, 1/8, 1/8 },

(5.2)

is C0 , shown in the margin.

C0 :

Using the extended code, we may encode acdbac as c+ (acdbac) = 100000100001010010000010.

(5.3)

There are basic requirements for a useful symbol code. First, any encoded string must have a unique decoding. Second, the symbol code must be easy to decode. And third, the code should achieve as much compression as possible.

Any encoded string must have a unique decoding A code C(X) is uniquely decodeable if, under the extended code C + , no two distinct strings have the same encoding, i.e., + + ∀ x, y ∈ A+ X , x 6= y ⇒ c (x) 6= c (y).

(5.4)

The code C0 defined above is an example of a uniquely decodeable code.

The symbol code must be easy to decode A symbol code is easiest to decode if it is possible to identify the end of a codeword as soon as it arrives, which means that no codeword can be a prefix of another codeword. [A word c is a prefix of another word d if there exists a tail string t such that the concatenation ct is identical to d. For example, 1 is a prefix of 101, and so is 10.] We will show later that we don’t lose any performance if we constrain our symbol code to be a prefix code. A symbol code is called a prefix code if no codeword is a prefix of any other codeword. A prefix code is also known as an instantaneous or self-punctuating code, because an encoded string can be decoded from left to right without looking ahead to subsequent codewords. The end of a codeword is immediately recognizable. A prefix code is uniquely decodeable. Prefix codes are also known as ‘prefix-free codes’ or ‘prefix condition codes’.

Prefix codes correspond to trees, as illustrated in the margin of the next page.

ai

c(ai )

li

a b c d

1000 0100 0010 0001

4 4 4 4

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5.1: Symbol codes

93

Example 5.4. The code C1 = {0, 101} is a prefix code because 0 is not a prefix of 101, nor is 101 a prefix of 0.

0

C1

1

0 0 1

101

Example 5.5. Let C2 = {1, 101}. This code is not a prefix code because 1 is a prefix of 101. Example 5.6. The code C3 = {0, 10, 110, 111} is a prefix code. Example 5.7. The code C4 = {00, 01, 10, 11} is a prefix code. Exercise 5.8.

[1, p.104]

0

0

C3 1

0

Is C2 uniquely decodeable?

1

Example 5.9. Consider exercise 4.1 (p.66) and figure 4.2 (p.69). Any weighing strategy that identifies the odd ball and whether it is heavy or light can be viewed as assigning a ternary code to each of the 24 possible states. This code is a prefix code.

The code should achieve as much compression as possible The expected length L(C, X) of a symbol code C for ensemble X is X L(C, X) = P (x) l(x). (5.5) x∈AX

10 0 110 1 111

0

C4 1

0

00

1 0

01 10

1

11

Prefix codes can be represented on binary trees. Complete prefix codes correspond to binary trees with no unused branches. C1 is an incomplete code.

We may also write this quantity as L(C, X) =

I X

pi li

(5.6)

i=1

where I = |AX |.

C3 :

Example 5.10. Let and

AX PX

= { a, b, c, d }, = { 1/2, 1/4, 1/8, 1/8 },

(5.7)

and consider the code C3 . The entropy of X is 1.75 bits, and the expected length L(C3 , X) of this code is also 1.75 bits. The sequence of symbols x = (acdbac) is encoded as c+ (x) = 0110111100110. C3 is a prefix code and is therefore uniquely decodeable. Notice that the codeword lengths satisfy li = log2 (1/pi ), or equivalently, pi = 2−li .

ai

c(ai )

pi

h(pi )

li

a b c d

0 10 110 111

1/2 1/4

1.0 2.0 3.0 3.0

1 2 3 3

Example 5.11. Consider the fixed length code for the same ensemble X, C 4 . The expected length L(C4 , X) is 2 bits.

a b c d

Example 5.12. Consider C5 . The expected length L(C5 , X) is 1.25 bits, which is less than H(X). But the code is not uniquely decodeable. The sequence x = (acdbac) encodes as 000111000, which can also be decoded as (cabdca). Example 5.13. Consider the code C6 . The expected length L(C6 , X) of this code is 1.75 bits. The sequence of symbols x = (acdbac) is encoded as c+ (x) = 0011111010011. Is C6 a prefix code? It is not, because c(a) = 0 is a prefix of both c(b) and c(c).

1/8 1/8

C4

C5

00 01 10 11

0 1 00 11

C6 : ai a b c d

c(ai )

pi

h(pi )

li

0 01 011 111

1/2

1.0 2.0 3.0 3.0

1 2 3 3

1/4 1/8 1/8

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94

5 — Symbol Codes Is C6 uniquely decodeable? This is not so obvious. If you think that it might not be uniquely decodeable, try to prove it so by finding a pair of strings x and y that have the same encoding. [The definition of unique decodeability is given in equation (5.4).] C6 certainly isn’t easy to decode. When we receive ‘00’, it is possible that x could start ‘aa’, ‘ab’ or ‘ac’. Once we have received ‘001111’, the second symbol is still ambiguous, as x could be ‘abd. . . ’ or ‘acd. . . ’. But eventually a unique decoding crystallizes, once the next 0 appears in the encoded stream. C6 is in fact uniquely decodeable. Comparing with the prefix code C 3 , we see that the codewords of C6 are the reverse of C3 ’s. That C3 is uniquely decodeable proves that C6 is too, since any string from C6 is identical to a string from C3 read backwards.

5.2 What limit is imposed by unique decodeability? We now ask, given a list of positive integers {l i }, does there exist a uniquely decodeable code with those integers as its codeword lengths? At this stage, we ignore the probabilities of the different symbols; once we understand unique decodeability better, we’ll reintroduce the probabilities and discuss how to make an optimal uniquely decodeable symbol code. In the examples above, we have observed that if we take a code such as {00, 01, 10, 11}, and shorten one of its codewords, for example 00 → 0, then we can retain unique decodeability only if we lengthen other codewords. Thus there seems to be a constrained budget that we can spend on codewords, with shorter codewords being more expensive. Let us explore the nature of this budget. If we build a code purely from codewords of length l equal to three, how many codewords can we have and retain unique decodeability? The answer is 2 l = 8. Once we have chosen all eight of these codewords, is there any way we could add to the code another codeword of some other length and retain unique decodeability? It would seem not. What if we make a code that includes a length-one codeword, ‘0’, with the other codewords being of length three? How many length-three codewords can we have? If we restrict attention to prefix codes, then we can have only four codewords of length three, namely {100, 101, 110, 111}. What about other codes? Is there any other way of choosing codewords of length 3 that can give more codewords? Intuitively, we think this unlikely. A codeword of length 3 appears to have a cost that is 22 times smaller than a codeword of length 1. Let’s define a total budget of size 1, which we can spend on codewords. If we set the cost of a codeword whose length is l to 2 −l , then we have a pricing system that fits the examples discussed above. Codewords of length 3 cost 1/8 each; codewords of length 1 cost 1/2 each. We can spend our budget on any codewords. If we go over our budget then the code will certainly not be uniquely decodeable. If, on the other hand, X i

2−li ≤ 1,

(5.8)

then the code may be uniquely decodeable. This inequality is the Kraft inequality. Kraft inequality. For any uniquely decodeable code C(X) over the binary

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5.2: What limit is imposed by unique decodeability?

95

alphabet {0, 1}, the codeword lengths must satisfy: I X i=1

2−li ≤ 1,

(5.9)

where I = |AX |. Completeness. If a uniquely decodeable code satisfies the Kraft inequality with equality then it is called a complete code. We want codes that are uniquely decodeable; prefix codes are uniquely decodeable, and are easy to decode. So life would be simpler for us if we could restrict attention to prefix codes. Fortunately, for any source there is an optimal symbol code that is also a prefix code. Kraft inequality and prefix codes. Given a set of codeword lengths that satisfy the Kraft inequality, there exists a uniquely decodeable prefix code with these codeword lengths. The Kraft inequality might be more accurately referred to as the Kraft– McMillan inequality: Kraft proved that if the inequality is satisfied, then a prefix code exists with the given lengths. McMillan (1956) proved the converse, that unique decodeability implies that the inequality holds.

Proof of the Kraft inequality. Define S = S

N

=

"

X i

2

−li

#N

=

I X I X

i1 =1 i2 =1

···

P

i

2−li . Consider the quantity

I X

2− (li1 + li2 + · · · liN ) .

(5.10)

iN =1

The quantity in the exponent, (li1 + li2 + · · · + liN ), is the length of the encoding of the string x = ai1 ai2 . . . aiN . For every string x of length N , there is one term in the above sum. Introduce an array A l that counts how many strings x have encoded length l. Then, defining l min = mini li and lmax = maxi li : NX lmax SN = 2−l Al . (5.11) l=N lmin

Now assume C is uniquely decodeable, so that for all x 6= y, c + (x) 6= c+ (y). Concentrate on the x that have encoded length l. There are a total of 2l distinct bit strings of length l, so it must be the case that Al ≤ 2l . So SN =

NX lmax

l=N lmin

2−l Al ≤

NX lmax

l=N lmin

1 ≤ N lmax .

(5.12)

Thus S N ≤ lmax N for all N . Now if S were greater than 1, then as N increases, S N would be an exponentially growing function, and for large enough N , an exponential always exceeds a polynomial such as l max N . But our result (S N ≤ lmax N ) is true for any N . Therefore S ≤ 1. 2 . Exercise 5.14.[3, p.104] Prove the result stated above, that for any set of codeword lengths {li } satisfying the Kraft inequality, there is a prefix code having those lengths.

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96

5 — Symbol Codes

Figure 5.1. The symbol coding budget. The ‘cost’ 2−l of each codeword (with length l) is indicated by the size of the box it is written in. The total budget available when making a uniquely decodeable code is 1. You can think of this diagram as showing a codeword supermarket, with the codewords arranged in aisles by their length, and the cost of each codeword indicated by the size of its box on the shelf. If the cost of the codewords that you take exceeds the budget then your code will not be uniquely decodeable.

0001

00

0010

001

0011

0

0100

010

0101

01

0110

011

0111 1000

100

1001

10

1010

101

1011

1

1100

110

1101

11

The total symbol code budget

0000

000

1110

111

1111

C0 000 00 001

   0000           0001                    0010        0011

0 010 01 011

            0100        0101 0110

     1000  0111

100 10 101 1 110

1001 1010 1011 1100 1101

11 111

 

  

  

  

  

  

  

  

  







  

  

  

  

  

  

  

  





                          



                         



                          



                         



                          



              0            



                          



                         



                          



                         



                          



        

1110 1111

1

C3

C4

000 00 001

0000 0001 0010 0011

010 01 011

                                                    100                                                        10                                                   101                  

   

 

 



 

 

 

 

 

 

            

   

 

 



  

 

 

 

 

 



   

 

 



  

110

   

 

 

  11



 



 



 



 



 



 



  

 

 

 

               

111       

             

 

 

 

 

 

 

 

    

0

0100 0101 0110 0111 1000 1001 1010 1011 1100 1101 1110 1111

1

                                00                                                                             01                                                                  10                                                          11                                  

000 001

0000 0001 0010 0011

010 011

0100 0101 0110 0111

100 101 110

C6 000 00 001

                                 01                                  

1000

1011 1100 1110 1111

                       011               

10

1010

101 1 110 11

0000 0001 0010 0011

010

100

1001

1101 111

                                                                       0                                                              

                       111               

0100 0101 0110 0111 1000 1001 1010 1011 1100 1101 1110 1111

Figure 5.2. Selections of codewords made by codes C0 , C3 , C4 and C6 from section 5.1.

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5.3: What’s the most compression that we can hope for?

97

A pictorial view of the Kraft inequality may help you solve this exercise. Imagine that we are choosing the codewords to make a symbol code. We can draw the set of all candidate codewords in a supermarket that displays the ‘cost’ of the codeword by the area of a box (figure 5.1). The total budget available – the ‘1’ on the right-hand side of the Kraft inequality – is shown at one side. Some of the codes discussed in section 5.1 are illustrated in figure 5.2. Notice that the codes that are prefix codes, C 0 , C3 , and C4 , have the property that to the right of any selected codeword, there are no other selected codewords – because prefix codes correspond to trees. Notice that a complete prefix code corresponds to a complete tree having no unused branches. We are now ready to put back the symbols’ probabilities {p i }. Given a set of symbol probabilities (the English language probabilities of figure 2.1, for example), how do we make the best symbol code – one with the smallest possible expected length L(C, X)? And what is that smallest possible expected length? It’s not obvious how to assign the codeword lengths. If we give short codewords to the more probable symbols then the expected length might be reduced; on the other hand, shortening some codewords necessarily causes others to lengthen, by the Kraft inequality.

5.3 What’s the most compression that we can hope for? We wish to minimize the expected length of a code, X L(C, X) = pi li .

(5.13)

i

As you might have guessed, the entropy appears as the lower bound on the expected length of a code. Lower bound on expected length. The expected length L(C, X) of a uniquely decodeable code is bounded below by H(X). P = i0 2−li0 , so Proof. We define the implicit probabilities q i ≡ 2−li /z, where zP that li = log 1/qi − log z. We then use Gibbs’ inequality, i pi log 1/qi ≥ P i pi log 1/pi , with equality if qi = pi , and the Kraft inequality z ≤ 1: X X pi log 1/qi − log z (5.14) pi li = L(C, X) = ≥

i

i

X

pi log 1/pi − log z

i

≥ H(X).

(5.15)

(5.16)

The equality L(C, X) = H(X) is achieved only if the Kraft equality z = 1 is satisfied, and if the codelengths satisfy l i = log(1/pi ). 2 This is an important result so let’s say it again: Optimal source codelengths. The expected length is minimized and is equal to H(X) only if the codelengths are equal to the Shannon information contents: li = log2 (1/pi ). (5.17) Implicit probabilities defined by codelengths. Conversely, any choice of codelengths {li } implicitly defines a probability distribution {q i }, qi ≡ 2−li /z,

(5.18)

for which those codelengths would be the optimal codelengths. If the code is complete then z = 1 and the implicit probabilities are given by qi = 2−li .

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98

5 — Symbol Codes

5.4 How much can we compress? So, we can’t compress below the entropy. How close can we expect to get to the entropy? Theorem 5.1 Source coding theorem for symbol codes. For an ensemble X there exists a prefix code C with expected length satisfying H(X) ≤ L(C, X) < H(X) + 1.

(5.19)

Proof. We set the codelengths to integers slightly larger than the optimum lengths: li = dlog 2 (1/pi )e (5.20)

where dl∗ e denotes the smallest integer greater than or equal to l ∗ . [We are not asserting that the optimal code necessarily uses these lengths, we are simply choosing these lengths because we can use them to prove the theorem.] We check that there is a prefix code with these lengths by confirming that the Kraft inequality is satisfied. X X X X pi = 1. (5.21) 2− log2 (1/pi ) = 2−dlog2 (1/pi )e ≤ 2−li = i

i

i

i

Then we confirm X X pi (log(1/pi ) + 1) = H(X) + 1. (5.22) pi dlog(1/pi )e < L(C, X) = i

i

2

The cost of using the wrong codelengths If we use a code whose lengths are not equal to the optimal codelengths, the average message length will be larger than the entropy. If the true probabilities are {pi } and we use a complete code with lengths li , we can view those lengths as defining implicit probabilities q i = 2−li . Continuing from equation (5.14), the average length is X pi log pi /qi , (5.23) L(C, X) = H(X) + i

i.e., it exceeds the entropy by the relative entropy D KL (p||q) (as defined on p.34).

5.5 Optimal source coding with symbol codes: Huffman coding Given a set of probabilities P, how can we design an optimal prefix code? For example, what is the best symbol code for the English language ensemble shown in figure 5.3? When we say ‘optimal’, let’s assume our aim is to minimize the expected length L(C, X).

How not to do it One might try to roughly split the set A X in two, and continue bisecting the subsets so as to define a binary tree from the root. This construction has the right spirit, as in the weighing problem, but it is not necessarily optimal; it achieves L(C, X) ≤ H(X) + 2.

x

P (x)

a b c d e f g h i j k l m n o p q r s t u v w x y z −

0.0575 0.0128 0.0263 0.0285 0.0913 0.0173 0.0133 0.0313 0.0599 0.0006 0.0084 0.0335 0.0235 0.0596 0.0689 0.0192 0.0008 0.0508 0.0567 0.0706 0.0334 0.0069 0.0119 0.0073 0.0164 0.0007 0.1928

Figure 5.3. An ensemble in need of a symbol code.

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5.5: Optimal source coding with symbol codes: Huffman coding

99

The Huffman coding algorithm We now present a beautifully simple algorithm for finding an optimal prefix code. The trick is to construct the code backwards starting from the tails of the codewords; we build the binary tree from its leaves. 1. Take the two least probable symbols in the alphabet. These two symbols will be given the longest codewords, which will have equal length, and differ only in the last digit.

Algorithm 5.4. Huffman coding algorithm.

2. Combine these two symbols into a single symbol, and repeat. Since each step reduces the size of the alphabet by one, this algorithm will have assigned strings to all the symbols after |A X | − 1 steps. Example 5.15. Let and x a b c d e

AX = { a, b, c, d, e } PX = { 0.25, 0.25, 0.2, 0.15, 0.15 }. step 1

step 2

step 3 step 4 0 0 0.25 0.25 0.25  0.55 1.0 0 0.25 0.25 0.45  0.45 1  0.2 0.2 1  0 0.15 0.3 0.3  1 0.15 1

The codewords are then obtained by concatenating the binary digits in reverse order: C = {00, 10, 11, 010, 011}. The codelengths selected by the Huffman algorithm (column 4 of table 5.5) are in some cases longer and in some cases shorter than the ideal codelengths, the Shannon information contents log 2 1/pi (column 3). The expected length of the code is L = 2.30 bits, whereas the entropy is H = 2.2855 bits. 2 If at any point there is more than one way of selecting the two least probable symbols then the choice may be made in any manner – the expected length of the code will not depend on the choice. Exercise 5.16.[3, p.105] Prove that there is no better symbol code for a source than the Huffman code. Example 5.17. We can make a Huffman code for the probability distribution over the alphabet introduced in figure 2.1. The result is shown in figure 5.6. This code has an expected length of 4.15 bits; the entropy of the ensemble is 4.11 bits. Observe the disparities between the assigned codelengths and the ideal codelengths log 2 1/pi .

Constructing a binary tree top-down is suboptimal In previous chapters we studied weighing problems in which we built ternary or binary trees. We noticed that balanced trees – ones in which, at every step, the two possible outcomes were as close as possible to equiprobable – appeared to describe the most efficient experiments. This gave an intuitive motivation for entropy as a measure of information content.

ai

pi

a b c d e

0.25 0.25 0.2 0.15 0.15

h(pi )

li

c(ai )

2.0 2.0 2.3 2.7 2.7

2 2 2 3 3

00 10 11 010 011

Table 5.5. Code created by the Huffman algorithm.

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100

5 — Symbol Codes

ai

pi

a b c d e f g h i j k l m n o p q r s t u v w x y z –

0.0575 0.0128 0.0263 0.0285 0.0913 0.0173 0.0133 0.0313 0.0599 0.0006 0.0084 0.0335 0.0235 0.0596 0.0689 0.0192 0.0008 0.0508 0.0567 0.0706 0.0334 0.0069 0.0119 0.0073 0.0164 0.0007 0.1928

1 pi

li

4.1 6.3 5.2 5.1 3.5 5.9 6.2 5.0 4.1 10.7 6.9 4.9 5.4 4.1 3.9 5.7 10.3 4.3 4.1 3.8 4.9 7.2 6.4 7.1 5.9 10.4 2.4

4 6 5 5 4 6 6 5 4 10 7 5 6 4 4 6 9 5 4 4 5 8 7 7 6 10 2

log2

c(ai ) 0000 001000 00101 10000 1100 111000 001001 10001 1001 1101000000 1010000 11101 110101 0001 1011 111001 110100001 11011 0011 1111 10101 11010001 1101001 1010001 101001 1101000001 01

Figure 5.6. Huffman code for the English language ensemble (monogram statistics).

a n b g c s − d h i k x y u o e j z q v w m r

f p

l t

It is not the case, however, that optimal codes can always be constructed by a greedy top-down method in which the alphabet is successively divided into subsets that are as near as possible to equiprobable. Example 5.18. Find the optimal binary symbol code for the ensemble: AX = { a, b, c, d, e, f, g } . PX = { 0.01, 0.24, 0.05, 0.20, 0.47, 0.01, 0.02 }

(5.24)

Notice that a greedy top-down method can split this set into two subsets {a, b, c, d} and {e, f, g} which both have probability 1/2, and that {a, b, c, d} can be divided into subsets {a, b} and {c, d}, which have probability 1/4; so a greedy top-down method gives the code shown in the third column of table 5.7, which has expected length 2.53. The Huffman coding algorithm yields the code shown in the fourth column, which has expected length 1.97. 2

5.6 Disadvantages of the Huffman code The Huffman algorithm produces an optimal symbol code for an ensemble, but this is not the end of the story. Both the word ‘ensemble’ and the phrase ‘symbol code’ need careful attention.

Changing ensemble If we wish to communicate a sequence of outcomes from one unchanging ensemble, then a Huffman code may be convenient. But often the appropriate

ai

pi

Greedy

Huffman

a b c d e f g

.01 .24 .05 .20 .47 .01 .02

000 001 010 011 10 110 111

000000 01 0001 001 1 000001 00001

Table 5.7. A greedily-constructed code compared with the Huffman code.

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5.6: Disadvantages of the Huffman code ensemble changes. If for example we are compressing text, then the symbol frequencies will vary with context: in English the letter u is much more probable after a q than after an e (figure 2.3). And furthermore, our knowledge of these context-dependent symbol frequencies will also change as we learn the statistical properties of the text source. Huffman codes do not handle changing ensemble probabilities with any elegance. One brute-force approach would be to recompute the Huffman code every time the probability over symbols changes. Another attitude is to deny the option of adaptation, and instead run through the entire file in advance and compute a good probability distribution, which will then remain fixed throughout transmission. The code itself must also be communicated in this scenario. Such a technique is not only cumbersome and restrictive, it is also suboptimal, since the initial message specifying the code and the document itself are partially redundant. This technique therefore wastes bits.

The extra bit An equally serious problem with Huffman codes is the innocuous-looking ‘extra bit’ relative to the ideal average length of H(X) – a Huffman code achieves a length that satisfies H(X) ≤ L(C, X) < H(X)+1, as proved in theorem 5.1. A Huffman code thus incurs an overhead of between 0 and 1 bits per symbol. If H(X) were large, then this overhead would be an unimportant fractional increase. But for many applications, the entropy may be as low as one bit per symbol, or even smaller, so the overhead L(C, X) − H(X) may dominate the encoded file length. Consider English text: in some contexts, long strings of characters may be highly predictable. For example, in the context ‘strings_of_ch’, one might predict the next nine symbols to be ‘aracters_’ with a probability of 0.99 each. A traditional Huffman code would be obliged to use at least one bit per character, making a total cost of nine bits where virtually no information is being conveyed (0.13 bits in total, to be precise). The entropy of English, given a good model, is about one bit per character (Shannon, 1948), so a Huffman code is likely to be highly inefficient. A traditional patch-up of Huffman codes uses them to compress blocks of symbols, for example the ‘extended sources’ X N we discussed in Chapter 4. The overhead per block is at most 1 bit so the overhead per symbol is at most 1/N bits. For sufficiently large blocks, the problem of the extra bit may be removed – but only at the expenses of (a) losing the elegant instantaneous decodeability of simple Huffman coding; and (b) having to compute the probabilities of all relevant strings and build the associated Huffman tree. One will end up explicitly computing the probabilities and codes for a huge number of strings, most of which will never actually occur. (See exercise 5.29 (p.103).)

Beyond symbol codes Huffman codes, therefore, although widely trumpeted as ‘optimal’, have many defects for practical purposes. They are optimal symbol codes, but for practical purposes we don’t want a symbol code. The defects of Huffman codes are rectified by arithmetic coding, which dispenses with the restriction that each symbol must translate into an integer number of bits. Arithmetic coding is the main topic of the next chapter.

101

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102

5 — Symbol Codes

5.7 Summary Kraft inequality. If a code is uniquely decodeable its lengths must satisfy X (5.25) 2−li ≤ 1. i

For any lengths satisfying the Kraft inequality, there exists a prefix code with those lengths. Optimal source codelengths for an ensemble are equal to the Shannon information contents 1 (5.26) li = log2 , pi and conversely, any choice of codelengths defines implicit probabilities qi =

2−li . z

(5.27)

The relative entropy DKL (p||q) measures how many bits per symbol are wasted by using a code whose implicit probabilities are q, when the ensemble’s true probability distribution is p. Source coding theorem for symbol codes. For an ensemble X, there exists a prefix code whose expected length satisfies H(X) ≤ L(C, X) < H(X) + 1.

(5.28)

The Huffman coding algorithm generates an optimal symbol code iteratively. At each iteration, the two least probable symbols are combined.

5.8 Exercises . Exercise 5.19.[2 ] Is the code {00, 11, 0101, 111, 1010, 100100, 0110} uniquely decodeable? . Exercise 5.20.[2 ] Is the ternary code {00, 012, 0110, 0112, 100, 201, 212, 22} uniquely decodeable? Exercise 5.21.[3, p.106] Make Huffman codes for X 2 , X 3 and X 4 where AX = {0, 1} and PX = {0.9, 0.1}. Compute their expected lengths and compare them with the entropies H(X 2 ), H(X 3 ) and H(X 4 ). Repeat this exercise for X 2 and X 4 where PX = {0.6, 0.4}. Exercise 5.22.[2, p.106] Find a probability distribution {p 1 , p2 , p3 , p4 } such that there are two optimal codes that assign different lengths {l i } to the four symbols. Exercise 5.23.[3 ] (Continuation of exercise 5.22.) Assume that the four probabilities {p1 , p2 , p3 , p4 } are ordered such that p1 ≥ p2 ≥ p3 ≥ p4 ≥ 0. Let Q be the set of all probability vectors p such that there are two optimal codes with different lengths. Give a complete description of Q. Find three probability vectors q(1) , q(2) , q(3) , which are the convex hull of Q, i.e., such that any p ∈ Q can be written as p = µ1 q(1) + µ2 q(2) + µ3 q(3) , where {µi } are positive.

(5.29)

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5.8: Exercises

103

. Exercise 5.24.[1 ] Write a short essay discussing how to play the game of twenty questions optimally. [In twenty questions, one player thinks of an object, and the other player has to guess the object using as few binary questions as possible, preferably fewer than twenty.] . Exercise 5.25.[2 ] Show that, if each probability pi is equal to an integer power of 2 then there exists a source code whose expected length equals the entropy. . Exercise 5.26.[2, p.106] Make ensembles for which the difference between the entropy and the expected length of the Huffman code is as big as possible. . Exercise 5.27.[2, p.106] A source X has an alphabet of eleven characters {a, b, c, d, e, f, g, h, i, j, k}, all of which have equal probability, 1/11. Find an optimal uniquely decodeable symbol code for this source. How much greater is the expected length of this optimal code than the entropy of X? . Exercise 5.28.[2 ] Consider the optimal symbol code for an ensemble X with alphabet size I from which all symbols have identical probability p = 1/I. I is not a power of 2. Show that the fraction f + of the I symbols that are assigned codelengths equal to l+ ≡ dlog 2 Ie (5.30) satisfies

+

2l I and that the expected length of the optimal symbol code is f+ = 2 −

L = l+ − 1 + f + .

(5.31)

(5.32)

By differentiating the excess length ∆L ≡ L − H(X) with respect to I, show that the excess length is bounded by ∆L ≤ 1 −

ln(ln 2) 1 − = 0.086. ln 2 ln 2

(5.33)

Exercise 5.29.[2 ] Consider a sparse binary source with P X = {0.99, 0.01}. Discuss how Huffman codes could be used to compress this source efficiently. Estimate how many codewords your proposed solutions require. . Exercise 5.30.[2 ] Scientific American carried the following puzzle in 1975. The poisoned glass. ‘Mathematicians are curious birds’, the police commissioner said to his wife. ‘You see, we had all those partly filled glasses lined up in rows on a table in the hotel kitchen. Only one contained poison, and we wanted to know which one before searching that glass for fingerprints. Our lab could test the liquid in each glass, but the tests take time and money, so we wanted to make as few of them as possible by simultaneously testing mixtures of small samples from groups of glasses. The university sent over a

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104

5 — Symbol Codes mathematics professor to help us. He counted the glasses, smiled and said: ‘ “Pick any glass you want, Commissioner. We’ll test it first.” ‘ “But won’t that waste a test?” I asked. ‘ “No,” he said, “it’s part of the best procedure. We can test one glass first. It doesn’t matter which one.” ’ ‘How many glasses were there to start with?’ the commissioner’s wife asked. ‘I don’t remember. Somewhere between 100 and 200.’ What was the exact number of glasses? Solve this puzzle and then explain why the professor was in fact wrong and the commissioner was right. What is in fact the optimal procedure for identifying the one poisoned glass? What is the expected waste relative to this optimum if one followed the professor’s strategy? Explain the relationship to symbol coding.

Exercise 5.31.[2, p.106] Assume that a sequence of symbols from the ensemble X introduced at the beginning of this chapter is compressed using the code C3 . Imagine picking one bit at random from the binary encoded sequence c = c(x1 )c(x2 )c(x3 ) . . . . What is the probability that this bit is a 1? . Exercise 5.32.[2, p.107] How should the binary Huffman encoding scheme be modified to make optimal symbol codes in an encoding alphabet with q symbols? (Also known as ‘radix q’.)

Mixture codes It is a tempting idea to construct a ‘metacode’ from several symbol codes that assign different-length codewords to the alternative symbols, then switch from one code to another, choosing whichever assigns the shortest codeword to the current symbol. Clearly we cannot do this for free. If one wishes to choose between two codes, then it is necessary to lengthen the message in a way that indicates which of the two codes is being used. If we indicate this choice by a single leading bit, it will be found that the resulting code is suboptimal because it is incomplete (that is, it fails the Kraft equality). Exercise 5.33.[3, p.108] Prove that this metacode is incomplete, and explain why this combined code is suboptimal.

5.9 Solutions Solution to exercise 5.8 (p.93). Yes, C 2 = {1, 101} is uniquely decodeable, even though it is not a prefix code, because no two different strings can map onto the same string; only the codeword c(a 2 ) = 101 contains the symbol 0. Solution to exercise 5.14 (p.95). We wish to prove that for any set of codeword lengths {li } satisfying the Kraft inequality, there is a prefix code having those lengths. This is readily proved by thinking of the codewords illustrated in figure 5.8 as being in a ‘codeword supermarket’, with size indicating cost. We imagine purchasing codewords one at a time, starting from the shortest codewords (i.e., the biggest purchases), using the budget shown at the right of figure 5.8. We start at one side of the codeword supermarket, say the

C3 : ai

c(ai )

pi

h(pi )

li

a b c d

0 10 110 111

1/2 1/4

1.0 2.0 3.0 3.0

1 2 3 3

1/8 1/8

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5.9: Solutions

105

00 001

0000 0001 0010 0011

0 010 01 011

0100 0101 0110 0111

100 10 101 1 110

1000 1001 1010 1011 1100 1101

11 111

The total symbol code budget

000

Figure 5.8. The codeword supermarket and the symbol coding budget. The ‘cost’ 2−l of each codeword (with length l) is indicated by the size of the box it is written in. The total budget available when making a uniquely decodeable code is 1.

1110 1111

symbol

probability

Huffman codewords

a

pa

cH (a)

b

pb

cH (b)

c

pc

cH (c)

Rival code’s codewords cR (a) cR (b) cR (c)

Modified rival code cR (c) cR (b) cR (a)

top, and purchase the first codeword of the required length. We advance down the supermarket a distance 2−l , and purchase the next codeword of the next required length, and so forth. Because the codeword lengths are getting longer, and the corresponding intervals are getting shorter, we can always buy an adjacent codeword to the latest purchase, so there is no wasting of PI −l the budget. Thus at the Ith we have advanced a distance i=1 2 i P codeword −l i down the supermarket; if 2 ≤ 1, we will have purchased all the codewords without running out of budget. Solution to exercise 5.16 (p.99). The proof that Huffman coding is optimal depends on proving that the key step in the algorithm – the decision to give the two symbols with smallest probability equal encoded lengths – cannot lead to a larger expected length than any other code. We can prove this by contradiction. Assume that the two symbols with smallest probability, called a and b, to which the Huffman algorithm would assign equal length codewords, do not have equal lengths in any optimal symbol code. The optimal symbol code is some other rival code in which these two codewords have unequal lengths la and lb with la < lb . Without loss of generality we can assume that this other code is a complete prefix code, because any codelengths of a uniquely decodeable code can be realized by a prefix code. In this rival code, there must be some other symbol c whose probability pc is greater than pa and whose length in the rival code is greater than or equal to lb , because the code for b must have an adjacent codeword of equal or greater length – a complete prefix code never has a solo codeword of the maximum length. Consider exchanging the codewords of a and c (figure 5.9), so that a is

Figure 5.9. Proof that Huffman coding makes an optimal symbol code. We assume that the rival code, which is said to be optimal, assigns unequal length codewords to the two symbols with smallest probability, a and b. By interchanging codewords a and c of the rival code, where c is a symbol with rival codelength as long as b’s, we can make a code better than the rival code. This shows that the rival code was not optimal.

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106

5 — Symbol Codes

encoded with the longer codeword that was c’s, and c, which is more probable than a, gets the shorter codeword. Clearly this reduces the expected length of the code. The change in expected length is (p a − pc )(lc − la ). Thus we have contradicted the assumption that the rival code is optimal. Therefore it is valid to give the two symbols with smallest probability equal encoded lengths. Huffman coding produces optimal symbol codes. 2 X2

Solution to exercise 5.21 (p.102). A Huffman code for where AX = {0, 1} and PX = {0.9, 0.1} is {00, 01, 10, 11} → {1, 01, 000, 001}. This code has L(C, X 2 ) = 1.29, whereas the entropy H(X 2 ) is 0.938. A Huffman code for X 3 is {000, 100, 010, 001, 101, 011, 110, 111} → {1, 011, 010, 001, 00000, 00001, 00010, 00011}. This has expected length L(C, X 3 ) = 1.598 whereas the entropy H(X 3 ) is 1.4069. A Huffman code for X 4 maps the sixteen source strings to the following codelengths: {0000, 1000, 0100, 0010, 0001, 1100, 0110, 0011, 0101, 1010, 1001, 1110, 1101, 1011, 0111, 1111} → {1, 3, 3, 3, 4, 6, 7, 7, 7, 7, 7, 9, 9, 9, 10, 10}. This has expected length L(C, X 4 ) = 1.9702 whereas the entropy H(X 4 ) is 1.876. When PX = {0.6, 0.4}, the Huffman code for X 2 has lengths {2, 2, 2, 2}; the expected length is 2 bits, and the entropy is 1.94 bits. A Huffman code for X 4 is shown in table 5.10. The expected length is 3.92 bits, and the entropy is 3.88 bits.

ai 0000 0001 0010 0100 1000 1100 1010 1001 0110 0101 0011 1110 1101 1011 0111 1111

pi

li

c(ai )

0.1296 0.0864 0.0864 0.0864 0.0864 0.0576 0.0576 0.0576 0.0576 0.0576 0.0576 0.0384 0.0384 0.0384 0.0384 0.0256

3 4 4 4 3 4 4 4 4 4 4 5 5 5 4 5

000 0100 0110 0111 100 1010 1100 1101 1110 1111 0010 00110 01010 01011 1011 00111

Table 5.10. Huffman code for X 4 when p0 = 0.6. Column 3 shows the assigned codelengths and column 4 the codewords. Some strings whose probabilities are identical, e.g., the fourth and fifth, receive different codelengths.

Solution to exercise 5.22 (p.102). The set of probabilities {p 1 , p2 , p3 , p4 } = {1/6, 1/6, 1/3, 1/3} gives rise to two different optimal sets of codelengths, because at the second step of the Huffman coding algorithm we can choose any of the three possible pairings. We may either put them in a constant length code {00, 01, 10, 11} or the code {000, 001, 01, 1}. Both codes have expected length 2. Another solution is {p1 , p2 , p3 , p4 } = {1/5, 1/5, 1/5, 2/5}. And a third is {p1 , p2 , p3 , p4 } = {1/3, 1/3, 1/3, 0}. Solution to exercise 5.26 (p.103). Let p max be the largest probability in p1 , p2 , . . . , pI . The difference between the expected length L and the entropy H can be no bigger than max(pmax , 0.086) (Gallager, 1978). See exercises 5.27–5.28 to understand where the curious 0.086 comes from. Solution to exercise 5.27 (p.103).

Length − entropy = 0.086.

Solution to exercise 5.31 (p.104). There are two ways to answer this problem correctly, and one popular way to answer it incorrectly. Let’s give the incorrect answer first: Erroneous answer. “We can pick a random bit by first picking a random source symbol xi with probability pi , then picking a random bit from c(xi ). If we define fi to be the fraction of the bits of c(xi ) that are 1s, we find X P (bit is 1) = pi fi (5.34) i

= 1/2 × 0 + 1/4 × 1/2 + 1/8 × 2/3 + 1/8 × 1 = 1/3.” (5.35)

C3 :

ai

c(ai )

pi

li

a b c d

0 10 110 111

1/2 1/4 1/8 1/8

1 2 3 3

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5.9: Solutions

107

This answer is wrong because it falls for the bus-stop fallacy, which was introduced in exercise 2.35 (p.38): if buses arrive at random, and we are interested in ‘the average time from one bus until the next’, we must distinguish two possible averages: (a) the average time from a randomly chosen bus until the next; (b) the average time between the bus you just missed and the next bus. The second ‘average’ is twice as big as the first because, by waiting for a bus at a random time, you bias your selection of a bus in favour of buses that follow a large gap. You’re unlikely to catch a bus that comes 10 seconds after a preceding bus! Similarly, the symbols c and d get encoded into longer-length binary strings than a, so when we pick a bit from the compressed string at random, we are more likely to land in a bit belonging to a c or a d than would be given by the probabilities pi in the expectation (5.34). All the probabilities need to be scaled up by li , and renormalized. Correct answer in the same style. Every time symbol x i is encoded, li bits are added to the binary string, of which f i li are 1s. The expected number of 1s added per symbol is X pi fi li ; (5.36) i

and the expected total number of bits added per symbol is X pi li .

(5.37)

So the fraction of 1s in the transmitted string is P pi fi li P (bit is 1) = Pi i pi li

(5.38)

i

=

1/2 × 0 + 1/4 × 1 + 1/8 × 2 + 1/8 × 3 7/4

7/8

= 7 = 1/2. /4

For a general symbol code and a general ensemble, the expectation (5.38) is the correct answer. But in this case, we can use a more powerful argument. Information-theoretic answer. The encoded string c is the output of an optimal compressor that compresses samples from X down to an expected length of H(X) bits. We can’t expect to compress this data any further. But if the probability P (bit is 1) were not equal to 1/2 then it would be possible to compress the binary string further (using a block compression code, say). Therefore P (bit is 1) must be equal to 1/2; indeed the probability of any sequence of l bits in the compressed stream taking on any particular value must be 2 −l . The output of a perfect compressor is always perfectly random bits. To put it another way, if the probability P (bit is 1) were not equal to then the information content per bit of the compressed string would be at most H2 (P (1)), which would be less than 1; but this contradicts the fact that we can recover the original data from c, so the information content per bit of the compressed string must be H(X)/L(C, X) = 1. 1/2,

Solution to exercise 5.32 (p.104). The general Huffman coding algorithm for an encoding alphabet with q symbols has one difference from the binary case. The process of combining q symbols into 1 symbol reduces the number of symbols by q − 1. So if we start with A symbols, we’ll only end up with a

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108

5 — Symbol Codes

complete q-ary tree if A mod (q −1) is equal to 1. Otherwise, we know that whatever prefix code we make, it must be an incomplete tree with a number of missing leaves equal, modulo (q −1), to A mod (q −1) − 1. For example, if a ternary tree is built for eight symbols, then there will unavoidably be one missing leaf in the tree. The optimal q-ary code is made by putting these extra leaves in the longest branch of the tree. This can be achieved by adding the appropriate number of symbols to the original source symbol set, all of these extra symbols having probability zero. The total number of leaves is then equal to r(q−1) + 1, for some integer r. The symbols are then repeatedly combined by taking the q symbols with smallest probability and replacing them by a single symbol, as in the binary Huffman coding algorithm. Solution to exercise 5.33 (p.104). We wish to show that a greedy metacode, which picks the code which gives the shortest encoding, is actually suboptimal, because it violates the Kraft inequality. We’ll assume that each symbol x is assigned lengths l k (x) by each of the candidate codes Ck . Let us assume there are K alternative codes and that we can encode which code is being used with a header of length log K bits. Then the metacode assigns lengths l 0 (x) that are given by l0 (x) = log2 K + min lk (x).

(5.39)

k

We compute the Kraft sum: X 1 X − mink lk (x) 0 2−l (x) = S= 2 . K x x

(5.40)

Let’s divide the set AX into non-overlapping subsets {Ak }K k=1 such that subset Ak contains all the symbols x that the metacode sends via code k. Then 1 X X −lk (x) S= 2 . (5.41) K k x∈Ak

Now if one sub-code k satisfies the Kraft equality must be the case that X 2−lk (x) ≤ 1,

P

x∈AX

2−lk (x) = 1, then it (5.42)

x∈Ak

with equality only if all the symbols x are in A k , which would mean that we are only using one of the K codes. So S≤

K 1 X 1 = 1, K

(5.43)

k=1

with equality only if equation (5.42) is an equality for all codes k. But it’s impossible for all the symbols to be in all the non-overlapping subsets {A k }K k=1 , so we can’t have equality (5.42) holding for all k. So S < 1. Another way of seeing that a mixture code is suboptimal is to consider the binary tree that it defines. Think of the special case of two codes. The first bit we send identifies which code we are using. Now, in a complete code, any subsequent binary string is a valid string. But once we know that we are using, say, code A, we know that what follows can only be a codeword corresponding to a symbol x whose encoding is shorter under code A than code B. So some strings are invalid continuations, and the mixture code is incomplete and suboptimal. For further discussion of this issue and its relationship to probabilistic modelling read about ‘bits back coding’ in section 28.3 and in Frey (1998).

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About Chapter 6 Before reading Chapter 6, you should have read the previous chapter and worked on most of the exercises in it. We’ll also make use of some Bayesian modelling ideas that arrived in the vicinity of exercise 2.8 (p.30).

109

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6 Stream Codes In this chapter we discuss two data compression schemes. Arithmetic coding is a beautiful method that goes hand in hand with the philosophy that compression of data from a source entails probabilistic modelling of that source. As of 1999, the best compression methods for text files use arithmetic coding, and several state-of-the-art image compression systems use it too. Lempel–Ziv coding is a ‘universal’ method, designed under the philosophy that we would like a single compression algorithm that will do a reasonable job for any source. In fact, for many real life sources, this algorithm’s universal properties hold only in the limit of unfeasibly large amounts of data, but, all the same, Lempel–Ziv compression is widely used and often effective.

6.1 The guessing game As a motivation for these two compression methods, consider the redundancy in a typical English text file. Such files have redundancy at several levels: for example, they contain the ASCII characters with non-equal frequency; certain consecutive pairs of letters are more probable than others; and entire words can be predicted given the context and a semantic understanding of the text. To illustrate the redundancy of English, and a curious way in which it could be compressed, we can imagine a guessing game in which an English speaker repeatedly attempts to predict the next character in a text file. For simplicity, let us assume that the allowed alphabet consists of the 26 upper case letters A,B,C,..., Z and a space ‘-’. The game involves asking the subject to guess the next character repeatedly, the only feedback being whether the guess is correct or not, until the character is correctly guessed. After a correct guess, we note the number of guesses that were made when the character was identified, and ask the subject to guess the next character in the same way. One sentence gave the following result when a human was asked to guess a sentence. The numbers of guesses are listed below each character. T H E R E - I S - N O - R E V E R S E - O N - A - M O T O R C Y C L E 1 1 1 5 1 1 2 1 1 2 1 1 15 1 17 1 1 1 2 1 3 2 1 2 2 7 1 1 1 1 4 1 1 1 1 1

Notice that in many cases, the next letter is guessed immediately, in one guess. In other cases, particularly at the start of syllables, more guesses are needed. What do this game and these results offer us? First, they demonstrate the redundancy of English from the point of view of an English speaker. Second, this game might be used in a data compression scheme, as follows. 110

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6.2: Arithmetic codes The string of numbers ‘1, 1, 1, 5, 1, . . . ’, listed above, was obtained by presenting the text to the subject. The maximum number of guesses that the subject will make for a given letter is twenty-seven, so what the subject is doing for us is performing a time-varying mapping of the twenty-seven letters {A, B, C, . . . , Z, −} onto the twenty-seven numbers {1, 2, 3, . . . , 27}, which we can view as symbols in a new alphabet. The total number of symbols has not been reduced, but since he uses some of these symbols much more frequently than others – for example, 1 and 2 – it should be easy to compress this new string of symbols. How would the uncompression of the sequence of numbers ‘1, 1, 1, 5, 1, . . . ’ work? At uncompression time, we do not have the original string ‘THERE. . . ’, we have only the encoded sequence. Imagine that our subject has an absolutely identical twin who also plays the guessing game with us, as if we knew the source text. If we stop him whenever he has made a number of guesses equal to the given number, then he will have just guessed the correct letter, and we can then say ‘yes, that’s right’, and move to the next character. Alternatively, if the identical twin is not available, we could design a compression system with the help of just one human as follows. We choose a window length L, that is, a number of characters of context to show the human. For every one of the 27L possible strings of length L, we ask them, ‘What would you predict is the next character?’, and ‘If that prediction were wrong, what would your next guesses be?’. After tabulating their answers to these 26 × 27 L questions, we could use two copies of these enormous tables at the encoder and the decoder in place of the two human twins. Such a language model is called an Lth order Markov model. These systems are clearly unrealistic for practical compression, but they illustrate several principles that we will make use of now.

6.2 Arithmetic codes When we discussed variable-length symbol codes, and the optimal Huffman algorithm for constructing them, we concluded by pointing out two practical and theoretical problems with Huffman codes (section 5.6). These defects are rectified by arithmetic codes, which were invented by Elias, by Rissanen and by Pasco, and subsequently made practical by Witten et al. (1987). In an arithmetic code, the probabilistic modelling is clearly separated from the encoding operation. The system is rather similar to the guessing game. The human predictor is replaced by a probabilistic model of the source. As each symbol is produced by the source, the probabilistic model supplies a predictive distribution over all possible values of the next symbol, that is, a list of positive numbers {p i } that sum to one. If we choose to model the source as producing i.i.d. symbols with some known distribution, then the predictive distribution is the same every time; but arithmetic coding can with equal ease handle complex adaptive models that produce context-dependent predictive distributions. The predictive model is usually implemented in a computer program. The encoder makes use of the model’s predictions to create a binary string. The decoder makes use of an identical twin of the model (just as in the guessing game) to interpret the binary string. Let the source alphabet be AX = {a1 , . . . , aI }, and let the Ith symbol aI have the special meaning ‘end of transmission’. The source spits out a sequence x1 , x2 , . . . , xn , . . . . The source does not necessarily produce i.i.d. symbols. We will assume that a computer program is provided to the encoder that assigns a

111

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112

6 — Stream Codes

predictive probability distribution over a i given the sequence that has occurred thus far, P (xn = ai | x1 , . . . , xn−1 ). The receiver has an identical program that produces the same predictive probability distribution P (x n = ai | x1 , . . . , xn−1 ). 0.00

Figure 6.1. Binary strings define real intervals within the real line [0,1). We first encountered a picture like this when we discussed the symbol-code supermarket in Chapter 5.

6

0.25

601 ?



0.50

0 ?

01101

6

0.75

1 ?

1.00

Concepts for understanding arithmetic coding Notation for intervals. The interval [0.01, 0.10) is all numbers between 0.01 and 0.10, including 0.010˙ ≡ 0.01000 . . . but not 0.100˙ ≡ 0.10000 . . . .

A binary transmission defines an interval within the real line from 0 to 1. For example, the string 01 is interpreted as a binary real number 0.01. . . , which corresponds to the interval [0.01, 0.10) in binary, i.e., the interval [0.25, 0.50) in base ten. The longer string 01101 corresponds to a smaller interval [0.01101, 0.01110). Because 01101 has the first string, 01, as a prefix, the new interval is a sub-interval of the interval [0.01, 0.10). A one-megabyte binary file (223 bits) is thus viewed as specifying a number between 0 and 1 to a precision of about two million decimal places – two million decimal digits, because each byte translates into a little more than two decimal digits. Now, we can also divide the real line [0,1) into I intervals of lengths equal to the probabilities P (x1 = ai ), as shown in figure 6.2. 0.00 P (x1 = a1 )

6 ?a1



6

a2



?

P (x1 = a1 ) + P (x1 = a2 ) .. . P (x1 = a1 ) + . . . + P (x1 = aI−1 ) 1.0

Figure 6.2. A probabilistic model defines real intervals within the real line [0,1).

a2 a1 a2 a5

.. . 6 ?aI

We may then take each interval ai and subdivide it into intervals denoted ai a1 , ai a2 , . . . , ai aI , such that the length of ai aj is proportional to P (x2 = aj | x1 = ai ). Indeed the length of the interval a i aj will be precisely the joint probability P (x1 = ai , x2 = aj ) = P (x1 = ai )P (x2 = aj | x1 = ai ).

(6.1)

Iterating this procedure, the interval [0, 1) can be divided into a sequence of intervals corresponding to all possible finite length strings x 1 x2 . . . xN , such that the length of an interval is equal to the probability of the string given our model.

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6.2: Arithmetic codes

113 Algorithm 6.3. Arithmetic coding. Iterative procedure to find the interval [u, v) for the string x1 x2 . . . x N .

u := 0.0 v := 1.0 p := v − u for n = 1 to N { Compute the cumulative probabilities Q n and Rn (6.2, 6.3) v := u + pRn (xn | x1 , . . . , xn−1 ) u := u + pQn (xn | x1 , . . . , xn−1 ) p := v − u }

Formulae describing arithmetic coding The process depicted in figure 6.2 can be written explicitly as follows. The intervals are defined in terms of the lower and upper cumulative probabilities Qn (ai | x1 , . . . , xn−1 )



Rn (ai | x1 , . . . , xn−1 )



i−1 X

=1 i X

i0

i0

=1

P (xn = ai0 | x1 , . . . , xn−1 ),

(6.2)

P (xn = ai0 | x1 , . . . , xn−1 ).

(6.3)

As the nth symbol arrives, we subdivide the n−1th interval at the points defined by Qn and Rn . For example, starting with the first symbol, the intervals ‘a1 ’, ‘a2 ’, and ‘aI ’ are a1 ↔ [Q1 (a1 ), R1 (a1 )) = [0, P (x1 = a1 )),

(6.4)

a2 ↔ [Q1 (a2 ), R1 (a2 )) = [P (x = a1 ), P (x = a1 ) + P (x = a2 )) ,

(6.5)

aI ↔ [Q1 (aI ), R1 (aI )) = [P (x1 = a1 ) + . . . + P (x1 = aI−1 ), 1.0) .

(6.6)

and

Algorithm 6.3 describes the general procedure.

To encode a string x1 x2 . . . xN , we locate the interval corresponding to x1 x2 . . . xN , and send a binary string whose interval lies within that interval. This encoding can be performed on the fly, as we now illustrate.

Example: compressing the tosses of a bent coin Imagine that we watch as a bent coin is tossed some number of times (cf. example 2.7 (p.30) and section 3.2 (p.51)). The two outcomes when the coin is tossed are denoted a and b. A third possibility is that the experiment is halted, an event denoted by the ‘end of file’ symbol, ‘2’. Because the coin is bent, we expect that the probabilities of the outcomes a and b are not equal, though beforehand we don’t know which is the more probable outcome. Encoding Let the source string be ‘bbba2’. We pass along the string one symbol at a time and use our model to compute the probability distribution of the next

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6 — Stream Codes

symbol given the string thus far. Let these probabilities be: Context (sequence thus far)

Probability of next symbol P (a) = 0.425

P (b) = 0.425

P (a | b) = 0.28

b

P (b | b) = 0.57

P (a | bb) = 0.21

bb

P (b | bb) = 0.64

P (a | bbb) = 0.17

bbb

P (b | bbb) = 0.68

P (a | bbba) = 0.28

bbba

P (b | bbba) = 0.57

P (2) = 0.15 P (2 | b) = 0.15

P (2 | bb) = 0.15

P (2 | bbb) = 0.15

P (2 | bbba) = 0.15

Figure 6.4 shows the corresponding intervals. The interval b is the middle 0.425 of [0, 1). The interval bb is the middle 0.567 of b, and so forth.

a

ba bba bbba b bb bbb bbbb bbb2 bb2 b2

2

00000 0000 00001 000 00010 0001 00011 00100 0010 00101 001 00110 0011 00111 01000 0100 01001 010 01010 0101 01011 01100 0110 01101 011 01110 0111 01111 10000 1000 10001 100 10010 1001 10011 10100 1010 10101 101 10110 1011 10111 11000 1100 11001 110 11010 1101 11011 11100 1110 11101 111 11110 1111 11111

Figure 6.4. Illustration of the arithmetic coding process as the sequence bbba2 is transmitted. 00

0

01

  10

 

bbbaa bbba

B B 1

B B

bbbab bbba2

10010111 10011000 10011001 10011010 10011011 10011100 10011101 10011110 OCC 10011111 C10100000 C 100111101

11

When the first symbol ‘b’ is observed, the encoder knows that the encoded string will start ‘01’, ‘10’, or ‘11’, but does not know which. The encoder writes nothing for the time being, and examines the next symbol, which is ‘b’. The interval ‘bb’ lies wholly within interval ‘1’, so the encoder can write the first bit: ‘1’. The third symbol ‘b’ narrows down the interval a little, but not quite enough for it to lie wholly within interval ‘10’. Only when the next ‘a’ is read from the source can we transmit some more bits. Interval ‘bbba’ lies wholly within the interval ‘1001’, so the encoder adds ‘001’ to the ‘1’ it has written. Finally when the ‘2’ arrives, we need a procedure for terminating the encoding. Magnifying the interval ‘bbba2’ (figure 6.4, right) we note that the marked interval ‘100111101’ is wholly contained by bbba2, so the encoding can be completed by appending ‘11101’.

10011

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6.2: Arithmetic codes Exercise 6.1.[2, p.127] Show that the overhead required to terminate a message is never more than 2 bits, relative to the ideal message length given the probabilistic model H, h(x | H) = log[1/P (x | H)]. This is an important result. Arithmetic coding is very nearly optimal. The message length is always within two bits of the Shannon information content of the entire source string, so the expected message length is within two bits of the entropy of the entire message. Decoding The decoder receives the string ‘100111101’ and passes along it one symbol at a time. First, the probabilities P (a), P (b), P (2) are computed using the identical program that the encoder used and the intervals ‘a’, ‘b’ and ‘2’ are deduced. Once the first two bits ‘10’ have been examined, it is certain that the original string must have been started with a ‘b’, since the interval ‘10’ lies wholly within interval ‘b’. The decoder can then use the model to compute P (a | b), P (b | b), P (2 | b) and deduce the boundaries of the intervals ‘ba’, ‘bb’ and ‘b2’. Continuing, we decode the second b once we reach ‘1001’, the third b once we reach ‘100111’, and so forth, with the unambiguous identification of ‘bbba2’ once the whole binary string has been read. With the convention that ‘2’ denotes the end of the message, the decoder knows to stop decoding. Transmission of multiple files How might one use arithmetic coding to communicate several distinct files over the binary channel? Once the 2 character has been transmitted, we imagine that the decoder is reset into its initial state. There is no transfer of the learnt statistics of the first file to the second file. If, however, we did believe that there is a relationship among the files that we are going to compress, we could define our alphabet differently, introducing a second end-of-file character that marks the end of the file but instructs the encoder and decoder to continue using the same probabilistic model.

The big picture Notice that to communicate a string of N letters both the encoder and the decoder needed to compute only N |A| conditional probabilities – the probabilities of each possible letter in each context actually encountered – just as in the guessing game. This cost can be contrasted with the alternative of using a Huffman code with a large block size (in order to reduce the possible onebit-per-symbol overhead discussed in section 5.6), where all block sequences that could occur must be considered and their probabilities evaluated. Notice how flexible arithmetic coding is: it can be used with any source alphabet and any encoded alphabet. The size of the source alphabet and the encoded alphabet can change with time. Arithmetic coding can be used with any probability distribution, which can change utterly from context to context. Furthermore, if we would like the symbols of the encoding alphabet (say, 0 and 1) to be used with unequal frequency, that can easily be arranged by subdividing the right-hand interval in proportion to the required frequencies.

How the probabilistic model might make its predictions The technique of arithmetic coding does not force one to produce the predictive probability in any particular way, but the predictive distributions might

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6 — Stream Codes

aaaa aaa aaab

aa

aab aaba aabb a aa2 abaa aba abab abba ab abb abbb ab2 a2 baaa baa baab baba ba bab babb ba2 bbaa bba bbab bbba b bb

bbb bbbb

bb2 b2

2

00000 00001 00010 00011 00100 00101 00110 00111 01000 01001 01010 01011 01100 01101 01110 01111 10000 10001 10010 10011 10100 10101 10110 10111 11000 11001 11010 11011 11100 11101 11110 11111

Figure 6.5. Illustration of the intervals defined by a simple Bayesian probabilistic model. The size of an intervals is proportional to the probability of the string. This model anticipates that the source is likely to be biased towards one of a and b, so sequences having lots of as or lots of bs have larger intervals than sequences of the same length that are 50:50 as and bs.

0000 000 0001 00 0010 001 0011 0 0100 010 0101 01 0110 011 0111 1000 100 1001 10 1010 101 1011 1 1100 110 1101 11 1110 111 1111

naturally be produced by a Bayesian model. Figure 6.4 was generated using a simple model that always assigns a probability of 0.15 to 2, and assigns the remaining 0.85 to a and b, divided in proportion to probabilities given by Laplace’s rule, PL (a | x1 , . . . , xn−1 ) =

Fa + 1 , Fa + F b + 2

(6.7)

where Fa (x1 , . . . , xn−1 ) is the number of times that a has occurred so far, and Fb is the count of bs. These predictions correspond to a simple Bayesian model that expects and adapts to a non-equal frequency of use of the source symbols a and b within a file. Figure 6.5 displays the intervals corresponding to a number of strings of length up to five. Note that if the string so far has contained a large number of bs then the probability of b relative to a is increased, and conversely if many as occur then as are made more probable. Larger intervals, remember, require fewer bits to encode.

Details of the Bayesian model Having emphasized that any model could be used – arithmetic coding is not wedded to any particular set of probabilities – let me explain the simple adaptive

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6.2: Arithmetic codes

117

probabilistic model used in the preceding example; we first encountered this model in exercise 2.8 (p.30). Assumptions The model will be described using parameters p2 , pa and pb , defined below, which should not be confused with the predictive probabilities in a particular context, for example, P (a | s = baa). A bent coin labelled a and b is tossed some number of times l, which we don’t know beforehand. The coin’s probability of coming up a when tossed is pa , and pb = 1 − pa ; the parameters pa , pb are not known beforehand. The source string s = baaba2 indicates that l was 5 and the sequence of outcomes was baaba. 1. It is assumed that the length of the string l has an exponential probability distribution P (l) = (1 − p2 )l p2 . (6.8) This distribution corresponds to assuming a constant probability p2 for the termination symbol ‘2’ at each character. 2. It is assumed that the non-terminal characters in the string are selected independently at random from an ensemble with probabilities P = {pa , pb }; the probability pa is fixed throughout the string to some unknown value that could be anywhere between 0 and 1. The probability of an a occurring as the next symbol, given pa (if only we knew it), is (1 − p2 )pa . The probability, given pa , that an unterminated string of length F is a given string s that contains {Fa , Fb } counts of the two outcomes is the Bernoulli distribution Fb a P (s | pa , F ) = pF (6.9) a (1 − pa ) . 3. We assume a uniform prior distribution for pa , P (pa ) = 1,

pa ∈ [0, 1],

(6.10)

and define pb ≡ 1 − pa . It would be easy to assume other priors on pa , with beta distributions being the most convenient to handle. This model was studied in section 3.2. The key result we require is the predictive distribution for the next symbol, given the string so far, s. This probability that the next character is a or b (assuming that it is not ‘2’) was derived in equation (3.16) and is precisely Laplace’s rule (6.7).

. Exercise 6.2.[3 ] Compare the expected message length when an ASCII file is compressed by the following three methods. Huffman-with-header. Read the whole file, find the empirical frequency of each symbol, construct a Huffman code for those frequencies, transmit the code by transmitting the lengths of the Huffman codewords, then transmit the file using the Huffman code. (The actual codewords don’t need to be transmitted, since we can use a deterministic method for building the tree given the codelengths.) Arithmetic code using the Laplace model. PL (a | x1 , . . . , xn−1 ) = P

Fa + 1 . a0 (Fa0 + 1)

(6.11)

Arithmetic code using a Dirichlet model. This model’s predictions are: Fa + α , (6.12) PD (a | x1 , . . . , xn−1 ) = P a0 (Fa0 + α)

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118

6 — Stream Codes where α is fixed to a number such as 0.01. A small value of α corresponds to a more responsive version of the Laplace model; the probability over characters is expected to be more nonuniform; α = 1 reproduces the Laplace model. Take care that the header of your Huffman message is self-delimiting. Special cases worth considering are (a) short files with just a few hundred characters; (b) large files in which some characters are never used.

6.3 Further applications of arithmetic coding Efficient generation of random samples Arithmetic coding not only offers a way to compress strings believed to come from a given model; it also offers a way to generate random strings from a model. Imagine sticking a pin into the unit interval at random, that line having been divided into subintervals in proportion to probabilities p i ; the probability that your pin will lie in interval i is p i . So to generate a sample from a model, all we need to do is feed ordinary random bits into an arithmetic decoder for that model. An infinite random bit sequence corresponds to the selection of a point at random from the line [0, 1), so the decoder will then select a string at random from the assumed distribution. This arithmetic method is guaranteed to use very nearly the smallest number of random bits possible to make the selection – an important point in communities where random numbers are expensive! [This is not a joke. Large amounts of money are spent on generating random bits in software and hardware. Random numbers are valuable.] A simple example of the use of this technique is in the generation of random bits with a nonuniform distribution {p 0 , p1 }. Exercise 6.3.[2, p.128] Compare the following two techniques for generating random symbols from a nonuniform distribution {p 0 , p1 } = {0.99, 0.01}: (a) The standard method: use a standard random number generator to generate an integer between 1 and 2 32 . Rescale the integer to (0, 1). Test whether this uniformly distributed random variable is less than 0.99, and emit a 0 or 1 accordingly. (b) Arithmetic coding using the correct model, fed with standard random bits. Roughly how many random bits will each method use to generate a thousand samples from this sparse distribution?

Efficient data-entry devices When we enter text into a computer, we make gestures of some sort – maybe we tap a keyboard, or scribble with a pointer, or click with a mouse; an efficient text entry system is one where the number of gestures required to enter a given text string is small. Writing can be viewed as an inverse process to data compression. In data compression, the aim is to map a given text string into a small number of bits. In text entry, we want a small sequence of gestures to produce our intended text. By inverting an arithmetic coder, we can obtain an information-efficient text entry device that is driven by continuous pointing gestures (Ward et al.,

Compression: text → bits Writing: text ←

gestures

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6.4: Lempel–Ziv coding

119

2000). In this system, called Dasher, the user zooms in on the unit interval to locate the interval corresponding to their intended string, in the same style as figure 6.4. A language model (exactly as used in text compression) controls the sizes of the intervals such that probable strings are quick and easy to identify. After an hour’s practice, a novice user can write with one finger driving Dasher at about 25 words per minute – that’s about half their normal ten-finger typing speed on a regular keyboard. It’s even possible to write at 25 words per minute, hands-free, using gaze direction to drive Dasher (Ward and MacKay, 2002). Dasher is available as free software for various platforms. 1

6.4 Lempel–Ziv coding The Lempel–Ziv algorithms, which are widely used for data compression (e.g., the compress and gzip commands), are different in philosophy to arithmetic coding. There is no separation between modelling and coding, and no opportunity for explicit modelling.

Basic Lempel–Ziv algorithm The method of compression is to replace a substring with a pointer to an earlier occurrence of the same substring. For example if the string is 1011010100010. . . , we parse it into an ordered dictionary of substrings that have not appeared before as follows: λ, 1, 0, 11, 01, 010, 00, 10, . . . . We include the empty substring λ as the first substring in the dictionary and order the substrings in the dictionary by the order in which they emerged from the source. After every comma, we look along the next part of the input sequence until we have read a substring that has not been marked off before. A moment’s reflection will confirm that this substring is longer by one bit than a substring that has occurred earlier in the dictionary. This means that we can encode each substring by giving a pointer to the earlier occurrence of that prefix and then sending the extra bit by which the new substring in the dictionary differs from the earlier substring. If, at the nth bit, we have enumerated s(n) substrings, then we can give the value of the pointer in dlog 2 s(n)e bits. The code for the above sequence is then as shown in the fourth line of the following table (with punctuation included for clarity), the upper lines indicating the source string and the value of s(n): source substrings λ s(n) 0 s(n)binary 000 (pointer, bit)

1 1 001 (, 1)

0 2 010 (0, 0)

11 3 011 (01, 1)

01 4 100 (10, 1)

010 5 101 (100, 0)

Notice that the first pointer we send is empty, because, given that there is only one substring in the dictionary – the string λ – no bits are needed to convey the ‘choice’ of that substring as the prefix. The encoded string is 100011101100001000010. The encoding, in this simple case, is actually a longer string than the source string, because there was no obvious redundancy in the source string. . Exercise 6.4.[2 ] Prove that any uniquely decodeable code from {0, 1} + to {0, 1}+ necessarily makes some strings longer if it makes some strings shorter. 1

http://www.inference.phy.cam.ac.uk/dasher/

00 6 110 (010, 0)

10 7 111 (001, 0)

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6 — Stream Codes

One reason why the algorithm described above lengthens a lot of strings is because it is inefficient – it transmits unnecessary bits; to put it another way, its code is not complete. Once a substring in the dictionary has been joined there by both of its children, then we can be sure that it will not be needed (except possibly as part of our protocol for terminating a message); so at that point we could drop it from our dictionary of substrings and shuffle them all along one, thereby reducing the length of subsequent pointer messages. Equivalently, we could write the second prefix into the dictionary at the point previously occupied by the parent. A second unnecessary overhead is the transmission of the new bit in these cases – the second time a prefix is used, we can be sure of the identity of the next bit. Decoding The decoder again involves an identical twin at the decoding end who constructs the dictionary of substrings as the data are decoded. . Exercise 6.5.[2, p.128] Encode the string 000000000000100000000000 using the basic Lempel–Ziv algorithm described above. . Exercise 6.6.[2, p.128] Decode the string 00101011101100100100011010101000011 that was encoded using the basic Lempel–Ziv algorithm. Practicalities In this description I have not discussed the method for terminating a string. There are many variations on the Lempel–Ziv algorithm, all exploiting the same idea but using different procedures for dictionary management, etc. The resulting programs are fast, but their performance on compression of English text, although useful, does not match the standards set in the arithmetic coding literature.

Theoretical properties In contrast to the block code, Huffman code, and arithmetic coding methods we discussed in the last three chapters, the Lempel–Ziv algorithm is defined without making any mention of a probabilistic model for the source. Yet, given any ergodic source (i.e., one that is memoryless on sufficiently long timescales), the Lempel–Ziv algorithm can be proven asymptotically to compress down to the entropy of the source. This is why it is called a ‘universal’ compression algorithm. For a proof of this property, see Cover and Thomas (1991). It achieves its compression, however, only by memorizing substrings that have happened so that it has a short name for them the next time they occur. The asymptotic timescale on which this universal performance is achieved may, for many sources, be unfeasibly long, because the number of typical substrings that need memorizing may be enormous. The useful performance of the algorithm in practice is a reflection of the fact that many files contain multiple repetitions of particular short sequences of characters, a form of redundancy to which the algorithm is well suited.

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6.5: Demonstration

Common ground I have emphasized the difference in philosophy behind arithmetic coding and Lempel–Ziv coding. There is common ground between them, though: in principle, one can design adaptive probabilistic models, and thence arithmetic codes, that are ‘universal’, that is, models that will asymptotically compress any source in some class to within some factor (preferably 1) of its entropy. However, for practical purposes, I think such universal models can only be constructed if the class of sources is severely restricted. A general purpose compressor that can discover the probability distribution of any source would be a general purpose artificial intelligence! A general purpose artificial intelligence does not yet exist.

6.5 Demonstration An interactive aid for exploring arithmetic coding, dasher.tcl, is available. 2 A demonstration arithmetic-coding software package written by Radford Neal3 consists of encoding and decoding modules to which the user adds a module defining the probabilistic model. It should be emphasized that there is no single general-purpose arithmetic-coding compressor; a new model has to be written for each type of source. Radford Neal’s package includes a simple adaptive model similar to the Bayesian model demonstrated in section 6.2. The results using this Laplace model should be viewed as a basic benchmark since it is the simplest possible probabilistic model – it simply assumes the characters in the file come independently from a fixed ensemble. The counts {Fi } of the symbols {ai } are rescaled and rounded as the file is read such that all the counts lie between 1 and 256. A state-of-the-art compressor for documents containing text and images, DjVu, uses arithmetic coding.4 It uses a carefully designed approximate arithmetic coder for binary alphabets called the Z-coder (Bottou et al., 1998), which is much faster than the arithmetic coding software described above. One of the neat tricks the Z-coder uses is this: the adaptive model adapts only occasionally (to save on computer time), with the decision about when to adapt being pseudo-randomly controlled by whether the arithmetic encoder emitted a bit. The JBIG image compression standard for binary images uses arithmetic coding with a context-dependent model, which adapts using a rule similar to Laplace’s rule. PPM (Teahan, 1995) is a leading method for text compression, and it uses arithmetic coding. There are many Lempel–Ziv-based programs. gzip is based on a version of Lempel–Ziv called ‘LZ77’ (Ziv and Lempel, 1977). compress is based on ‘LZW’ (Welch, 1984). In my experience the best is gzip, with compress being inferior on most files. bzip is a block-sorting file compressor, which makes use of a neat hack called the Burrows–Wheeler transform (Burrows and Wheeler, 1994). This method is not based on an explicit probabilistic model, and it only works well for files larger than several thousand characters; but in practice it is a very effective compressor for files in which the context of a character is a good predictor for that character.5 2

http://www.inference.phy.cam.ac.uk/mackay/itprnn/softwareI.html ftp://ftp.cs.toronto.edu/pub/radford/www/ac.software.html 4 http://www.djvuzone.org/ 5 There is a lot of information about the Burrows–Wheeler transform on the net. http://dogma.net/DataCompression/BWT.shtml 3

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122

6 — Stream Codes

Compression of a text file Table 6.6 gives the computer time in seconds taken and the compression achieved when these programs are applied to the LATEX file containing the text of this chapter, of size 20,942 bytes. Method Laplace model gzip compress

Compression time / sec

Compressed size (%age of 20,942)

Uncompression time / sec

0.28 0.10 0.05

12 974 (61%) 8 177 (39%) 10 816 (51%)

0.32 0.01 0.05

Table 6.6. Comparison of compression algorithms applied to a text file.

7 495 (36%) 7 640 (36%) 6 800 (32%)

bzip bzip2 ppmz

Compression of a sparse file Interestingly, gzip does not always do so well. Table 6.7 gives the compression achieved when these programs are applied to a text file containing 10 6 characters, each of which is either 0 and 1 with probabilities 0.99 and 0.01. The Laplace model is quite well matched to this source, and the benchmark arithmetic coder gives good performance, followed closely by compress; gzip is worst. An ideal model for this source would compress the file into about 106 H2 (0.01)/8 ' 10 100 bytes. The Laplace-model compressor falls short of this performance because it is implemented using only eight-bit precision. The ppmz compressor compresses the best of all, but takes much more computer time. Method

Compression time / sec

Laplace model gzip gzip --best+ compress

0.45 0.22 1.63 0.13

bzip bzip2 ppmz

0.30 0.19 533

Compressed size / bytes 14 143 20 646 15 553 14 785

(1.4%) (2.1%) (1.6%) (1.5%)

10 903 (1.09%) 11 260 (1.12%) 10 447 (1.04%)

Uncompression time / sec 0.57 0.04 0.05 0.03 0.17 0.05 535

6.6 Summary In the last three chapters we have studied three classes of data compression codes. Fixed-length block codes (Chapter 4). These are mappings from a fixed number of source symbols to a fixed-length binary message. Only a tiny fraction of the source strings are given an encoding. These codes were fun for identifying the entropy as the measure of compressibility but they are of little practical use.

Table 6.7. Comparison of compression algorithms applied to a random file of 106 characters, 99% 0s and 1% 1s.

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6.7: Exercises on stream codes Symbol codes (Chapter 5). Symbol codes employ a variable-length code for each symbol in the source alphabet, the codelengths being integer lengths determined by the probabilities of the symbols. Huffman’s algorithm constructs an optimal symbol code for a given set of symbol probabilities. Every source string has a uniquely decodeable encoding, and if the source symbols come from the assumed distribution then the symbol code will compress to an expected length per character L lying in the interval [H, H + 1). Statistical fluctuations in the source may make the actual length longer or shorter than this mean length. If the source is not well matched to the assumed distribution then the mean length is increased by the relative entropy D KL between the source distribution and the code’s implicit distribution. For sources with small entropy, the symbol has to emit at least one bit per source symbol; compression below one bit per source symbol can be achieved only by the cumbersome procedure of putting the source data into blocks. Stream codes. The distinctive property of stream codes, compared with symbol codes, is that they are not constrained to emit at least one bit for every symbol read from the source stream. So large numbers of source symbols may be coded into a smaller number of bits. This property could be obtained using a symbol code only if the source stream were somehow chopped into blocks. • Arithmetic codes combine a probabilistic model with an encoding algorithm that identifies each string with a sub-interval of [0, 1) of size equal to the probability of that string under the model. This code is almost optimal in the sense that the compressed length of a string x closely matches the Shannon information content of x given the probabilistic model. Arithmetic codes fit with the philosophy that good compression requires data modelling, in the form of an adaptive Bayesian model. • Lempel–Ziv codes are adaptive in the sense that they memorize strings that have already occurred. They are built on the philosophy that we don’t know anything at all about what the probability distribution of the source will be, and we want a compression algorithm that will perform reasonably well whatever that distribution is. Both arithmetic codes and Lempel–Ziv codes will fail to decode correctly if any of the bits of the compressed file are altered. So if compressed files are to be stored or transmitted over noisy media, error-correcting codes will be essential. Reliable communication over unreliable channels is the topic of Part II.

6.7 Exercises on stream codes Exercise 6.7.[2 ] Describe an arithmetic coding algorithm to encode random bit strings of length N and weight K (i.e., K ones and N − K zeroes) where N and K are given. For the case N = 5, K = 2, show in detail the intervals corresponding to all source substrings of lengths 1–5. . Exercise 6.8.[2, p.128] How many bits are needed to specify a selection of K objects from N objects? (N and K are assumed to be known and the

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124

6 — Stream Codes selection of K objects is unordered.) How might such a selection be made at random without being wasteful of random bits?

. Exercise 6.9.[2 ] A binary source X emits independent identically distributed symbols with probability distribution {f 0 , f1 }, where f1 = 0.01. Find an optimal uniquely-decodeable symbol code for a string x = x 1 x2 x3 of three successive samples from this source. Estimate (to one decimal place) the factor by which the expected length of this optimal code is greater than the entropy of the three-bit string x. [H2 (0.01) ' 0.08, where H2 (x) = x log 2 (1/x) + (1 − x) log 2 (1/(1 − x)).]

An arithmetic code is used to compress a string of 1000 samples from the source X. Estimate the mean and standard deviation of the length of the compressed file.

. Exercise 6.10.[2 ] Describe an arithmetic coding algorithm to generate random bit strings of length N with density f (i.e., each bit has probability f of being a one) where N is given. Exercise 6.11.[2 ] Use a modified Lempel–Ziv algorithm in which, as discussed on p.120, the dictionary of prefixes is pruned by writing new prefixes into the space occupied by prefixes that will not be needed again. Such prefixes can be identified when both their children have been added to the dictionary of prefixes. (You may neglect the issue of termination of encoding.) Use this algorithm to encode the string 0100001000100010101000001. Highlight the bits that follow a prefix on the second occasion that that prefix is used. (As discussed earlier, these bits could be omitted.) Exercise 6.12.[2, p.128] Show that this modified Lempel–Ziv code is still not ‘complete’, that is, there are binary strings that are not encodings of any string. . Exercise 6.13.[3, p.128] Give examples of simple sources that have low entropy but would not be compressed well by the Lempel–Ziv algorithm.

6.8 Further exercises on data compression The following exercises may be skipped by the reader who is eager to learn about noisy channels. Exercise 6.14.[3, p.130] Consider a Gaussian distribution in N dimensions,  P 2 x 1 (6.13) P (x) = exp − n 2 n . 2 N/2 2σ (2πσ ) P 2 1/2 the mean Define the radius of a point x to be r = n xn P . Estimate  2 and variance of the square of the radius, r 2 = n xn . You may find helpful the integral   Z 1 x2 4 dx x exp − = 3σ 4 , 2σ 2 (2πσ 2 )1/2

(6.14)

though you should be able to estimate the required quantities without it.

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6.8: Further exercises on data compression

125

Assuming that N is large, show that nearly√all the probability of a Gaussian is contained in a thin shell of radius N σ. Find the thickness of the shell. Evaluate the probability density (6.13) at a point in that thin shell and at the origin x = 0 and compare. Use the case N = 1000 as an example. Notice that nearly all the probability mass is located in a different part of the space from the region of highest probability density. Exercise 6.15.[2 ] Explain what is meant by an optimal binary symbol code. Find an optimal binary symbol code for the ensemble: A = {a, b, c, d, e, f, g, h, i, j}, P=



2 4 5 6 8 9 10 25 30 1 , , , , , , , , , 100 100 100 100 100 100 100 100 100 100



,

and compute the expected length of the code. Exercise 6.16.[2 ] A string y = x1 x2 consists of two independent samples from an ensemble   1 3 6 , , . X : AX = {a, b, c}; PX = 10 10 10 What is the entropy of y? Construct an optimal binary symbol code for the string y, and find its expected length. Exercise 6.17.[2 ] Strings of N independent samples from an ensemble with P = {0.1, 0.9} are compressed using an arithmetic code that is matched to that ensemble. Estimate the mean and standard deviation of the compressed strings’ lengths for the case N = 1000. [H 2 (0.1) ' 0.47] Exercise 6.18.[3 ] Source coding with variable-length symbols. In the chapters on source coding, we assumed that we were encoding into a binary alphabet {0, 1} in which both symbols should be used with equal frequency. In this question we explore how the encoding alphabet should be used if the symbols take different times to transmit. A poverty-stricken student communicates for free with a friend using a telephone by selecting an integer n ∈ {1, 2, 3 . . .}, making the friend’s phone ring n times, then hanging up in the middle of the nth ring. This process is repeated so that a string of symbols n 1 n2 n3 . . . is received. What is the optimal way to communicate? If large integers n are selected then the message takes longer to communicate. If only small integers n are used then the information content per symbol is small. We aim to maximize the rate of information transfer, per unit time. Assume that the time taken to transmit a number of rings n and to redial is ln seconds. Consider a probability distribution over n, {p n }. Defining the average duration per symbol to be X L(p) = pn ln (6.15) n

probability density is maximized here √

Nσ almost all probability mass is here

Figure 6.8. Schematic representation of the typical set of an N -dimensional Gaussian distribution.

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126

6 — Stream Codes and the entropy per symbol to be H(p) =

X

pn log2

n

1 , pn

(6.16)

show that for the average information rate per second to be maximized, the symbols must be used with probabilities of the form pn = where Z =

P

n2

−βln

1 −βln 2 Z

(6.17)

and β satisfies the implicit equation β=

H(p) , L(p)

(6.18)

that is, β is the rate of communication. Show that these two equations (6.17, 6.18) imply that β must be set such that log Z = 0.

(6.19)

Assuming that the channel has the property ln = n seconds,

(6.20)

find the optimal distribution p and show that the maximal information rate is 1 bit per second. How does this compare with the information rate per second achieved if p is set to (1/2, 1/2, 0, 0, 0, 0, . . .) — that is, only the symbols n = 1 and n = 2 are selected, and they have equal probability? Discuss the relationship between the results (6.17, 6.19) derived above, and the Kraft inequality from source coding theory. How might a random binary source be efficiently encoded into a sequence of symbols n1 n2 n3 . . . for transmission over the channel defined in equation (6.20)? . Exercise 6.19.[1 ] How many bits does it take to shuffle a pack of cards? . Exercise 6.20.[2 ] In the card game Bridge, the four players receive 13 cards each from the deck of 52 and start each game by looking at their own hand and bidding. The legal bids are, in ascending order 1♣, 1♦, 1♥, 1♠, 1N T, 2♣, 2♦, . . . 7♥, 7♠, 7N T , and successive bids must follow this order; a bid of, say, 2♥ may only be followed by higher bids such as 2♠ or 3♣ or 7N T . (Let us neglect the ‘double’ bid.) The players have several aims when bidding. One of the aims is for two partners to communicate to each other as much as possible about what cards are in their hands. Let us concentrate on this task. (a) After the cards have been dealt, how many bits are needed for North to convey to South what her hand is? (b) Assuming that E and W do not bid at all, what is the maximum total information that N and S can convey to each other while bidding? Assume that N starts the bidding, and that once either N or S stops bidding, the bidding stops.

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6.9: Solutions

127

. Exercise 6.21.[2 ] My old ‘arabic’ microwave oven had 11 buttons for entering cooking times, and my new ‘roman’ microwave has just five. The buttons of the roman microwave are labelled ‘10 minutes’, ‘1 minute’, ‘10 seconds’, ‘1 second’, and ‘Start’; I’ll abbreviate these five strings to the symbols M, C, X, I, 2. To enter one minute and twenty-three seconds (1:23), the arabic sequence is 1232,

(6.21)

CXXIII2.

(6.22)

and the roman sequence is

Each of these keypads defines a code mapping the 3599 cooking times from 0:01 to 59:59 into a string of symbols. (a) Which times can be produced with two or three symbols? (For example, 0:20 can be produced by three symbols in either code: XX2 and 202.) (b) Are the two codes complete? Give a detailed answer. (c) For each code, name a cooking time that it can produce in four symbols that the other code cannot. (d) Discuss the implicit probability distributions over times to which each of these codes is best matched. (e) Concoct a plausible probability distribution over times that a real user might use, and evaluate roughly the expected number of symbols, and maximum number of symbols, that each code requires. Discuss the ways in which each code is inefficient or efficient. (f) Invent a more efficient cooking-time-encoding system for a microwave oven. Exercise 6.22.[2, p.132] Is the standard binary representation for positive integers (e.g. cb (5) = 101) a uniquely decodeable code? Design a binary code for the positive integers, i.e., a mapping from n ∈ {1, 2, 3, . . .} to c(n) ∈ {0, 1}+ , that is uniquely decodeable. Try to codes that are prefix codes and that satisfy the Kraft equality P design −ln = 1. 2 n

Motivations: any data file terminated by a special end of file character can be mapped onto an integer, so a prefix code for integers can be used as a self-delimiting encoding of files too. Large files correspond to large integers. Also, one of the building blocks of a ‘universal’ coding scheme – that is, a coding scheme that will work OK for a large variety of sources – is the ability to encode integers. Finally, in microwave ovens, cooking times are positive integers!

Discuss criteria by which one might compare alternative codes for integers (or, equivalently, alternative self-delimiting codes for files).

6.9 Solutions Solution to exercise 6.1 (p.115). The worst-case situation is when the interval to be represented lies just inside a binary interval. In this case, we may choose either of two binary intervals as shown in figure 6.10. These binary intervals

Arabic 1 4 7

2 5 8 0

Roman 3 6 9 2

M C

X I

2

Figure 6.9. Alternative keypads for microwave ovens.

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128

6 — Stream Codes

Source string’s interval

Figure 6.10. Termination of arithmetic coding in the worst case, where there is a two bit overhead. Either of the two binary intervals marked on the right-hand side may be chosen. These binary intervals are no smaller than P (x|H)/4.

Binary intervals

6 6 P (x|H)

? 6 ?

?

are no smaller than P (x|H)/4, so the binary encoding has a length no greater than log 2 1/P (x|H) + log 2 4, which is two bits more than the ideal message length. Solution to exercise 6.3 (p.118). The standard method uses 32 random bits per generated symbol and so requires 32 000 bits to generate one thousand samples. Arithmetic coding uses on average about H 2 (0.01) = 0.081 bits per generated symbol, and so requires about 83 bits to generate one thousand samples (assuming an overhead of roughly two bits associated with termination). Fluctuations in the number of 1s would produce variations around this mean with standard deviation 21. Solution to exercise 6.5 (p.120). which comes from the parsing

The encoding is 010100110010110001100,

0, 00, 000, 0000, 001, 00000, 000000

(6.23)

which is encoded thus: (, 0), (1, 0), (10, 0), (11, 0), (010, 1), (100, 0), (110, 0).

(6.24)

Solution to exercise 6.6 (p.120). The decoding is 0100001000100010101000001. Solution to exercise 6.8 (p.123). This problem is equivalent to exercise 6.7 (p.123).  The selection of K objects from N objects requires dlog 2 N K e bits ' N H2 (K/N ) bits. This selection could be made using arithmetic coding. The selection corresponds to a binary string of length N in which the 1 bits represent which objects are selected. Initially the probability of a 1 is K/N and the probability of a 0 is (N −K)/N . Thereafter, given that the emitted string thus far, of length n, contains k 1s, the probability of a 1 is (K −k)/(N −n) and the probability of a 0 is 1 − (K −k)/(N −n). Solution to exercise 6.12 (p.124). This modified Lempel–Ziv code is still not ‘complete’, because, for example, after five prefixes have been collected, the pointer could be any of the strings 000, 001, 010, 011, 100, but it cannot be 101, 110 or 111. Thus there are some binary strings that cannot be produced as encodings. Solution to exercise 6.13 (p.124). Sources with low entropy that are not well compressed by Lempel–Ziv include:

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6.9: Solutions (a) Sources with some symbols that have long range correlations and intervening random junk. An ideal model should capture what’s correlated and compress it. Lempel–Ziv can compress the correlated features only by memorizing all cases of the intervening junk. As a simple example, consider a telephone book in which every line contains an (old number, new number) pair: 285-3820:572-58922 258-8302:593-20102 The number of characters per line is 18, drawn from the 13-character alphabet {0, 1, . . . , 9, −, :, 2}. The characters ‘-’, ‘:’ and ‘2’ occur in a predictable sequence, so the true information content per line, assuming all the phone numbers are seven digits long, and assuming that they are random sequences, is about 14 bans. (A ban is the information content of a random integer between 0 and 9.) A finite state language model could easily capture the regularities in these data. A Lempel–Ziv algorithm will take a long time before it compresses such a file down to 14 bans per line, however, because in order for it to ‘learn’ that the string :ddd is always followed by -, for any three digits ddd, it will have to see all those strings. So near-optimal compression will only be achieved after thousands of lines of the file have been read.

Figure 6.11. A source with low entropy that is not well compressed by Lempel–Ziv. The bit sequence is read from left to right. Each line differs from the line above in f = 5% of its bits. The image width is 400 pixels.

(b) Sources with long range correlations, for example two-dimensional images that are represented by a sequence of pixels, row by row, so that vertically adjacent pixels are a distance w apart in the source stream, where w is the image width. Consider, for example, a fax transmission in which each line is very similar to the previous line (figure 6.11). The true entropy is only H2 (f ) per pixel, where f is the probability that a pixel differs from its parent. Lempel–Ziv algorithms will only compress down to the entropy once all strings of length 2 w = 2400 have occurred and their successors have been memorized. There are only about 2 300 particles in the universe, so we can confidently say that Lempel–Ziv codes will never capture the redundancy of such an image. Another highly redundant texture is shown in figure 6.12. The image was made by dropping horizontal and vertical pins randomly on the plane. It contains both long-range vertical correlations and long-range horizontal correlations. There is no practical way that Lempel–Ziv, fed with a pixel-by-pixel scan of this image, could capture both these correlations. Biological computational systems can readily identify the redundancy in these images and in images that are much more complex; thus we might anticipate that the best data compression algorithms will result from the development of artificial intelligence methods.

129

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130

6 — Stream Codes

Figure 6.12. A texture consisting of horizontal and vertical pins dropped at random on the plane.

(c) Sources with intricate redundancy, such as files generated by computers. For example, a LATEX file followed by its encoding into a PostScript file. The information content of this pair of files is roughly equal to the information content of the LATEX file alone. (d) A picture of the Mandelbrot set. The picture has an information content equal to the number of bits required to specify the range of the complex plane studied, the pixel sizes, and the colouring rule used. (e) A picture of a ground state of a frustrated antiferromagnetic Ising model (figure 6.13), which we will discuss in Chapter 31. Like figure 6.12, this binary image has interesting correlations in two directions. Figure 6.13. Frustrated triangular Ising model in one of its ground states.

(f) Cellular automata – figure 6.14 shows the state history of 100 steps of a cellular automaton with 400 cells. The update rule, in which each cell’s new state depends on the state of five preceding cells, was selected at random. The information content is equal to the information in the boundary (400 bits), and the propagation rule, which here can be described in 32 bits. An optimal compressor will thus give a compressed file length which is essentially constant, independent of the vertical height of the image. Lempel–Ziv would only give this zero-cost compression once the cellular automaton has entered a periodic limit cycle, which could easily take about 2100 iterations. In contrast, the JBIG compression method, which models the probability of a pixel given its local context and uses arithmetic coding, would do a good job on these images. Solution to exercise 6.14 (p.124). For a one-dimensional Gaussian, the variance of x, E[x2 ], is σ 2 . So the mean value of r 2 in N dimensions, since the components of x are independent random variables, is E[r 2 ] = N σ 2 .

(6.25)

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6.9: Solutions

131

Figure 6.14. The 100-step time-history of a cellular automaton with 400 cells.

The variance of r 2 , similarly, is N times the variance of x 2 , where x is a one-dimensional Gaussian variable.   Z x2 1 4 2 x exp − 2 − σ 4 . (6.26) var(x ) = dx 2σ (2πσ 2 )1/2 The integral is found to be 3σ 4 (equation (6.14)), so var(x2 ) = 2σ 4 . Thus the variance of r 2 is 2N σ 4 . 2 For large N , the central-limit theorem indicates √ that 2r has a Gaussian 2 distribution with mean N σ and standard deviation 2N √σ , so the probability density of r must similarly be concentrated about r ' N σ. The thickness of this shell is given by turning the standard deviation of r 2 into a standard deviation on r: for small √ δr/r, δ log r = δr/r = (1/2)δ log r 2 = (1/2)δ(r 2 )/r 2 , so √ setting δ(r 2 ) = 2N σ 2 , r has standard deviation δr = (1/2)rδ(r 2 )/r 2 = σ/ 2. √ The probability density of the Gaussian at a point x shell where r = N σ is     1 N σ2 N 1 exp − 2 = exp − P (xshell ) = . (6.27) 2σ 2 (2πσ 2 )N/2 (2πσ 2 )N/2 Whereas the probability density at the origin is P (x = 0) =

1 . (2πσ 2 )N/2

(6.28)

Thus P (xshell )/P (x = 0) = exp (−N/2) . The probability density at the typical radius is e−N/2 times smaller than the density at the origin. If N = 1000, then the probability density at the origin is e 500 times greater.

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7 Codes for Integers

This chapter is an aside, which may safely be skipped.

Solution to exercise 6.22 (p.127) To discuss the coding of integers we need some definitions. The standard binary representation of a positive integer n will be denoted by cb (n), e.g., cb (5) = 101, cb (45) = 101101. The standard binary length of a positive integer n, l b (n), length of the string cb (n). For example, lb (5) = 3, lb (45) = 6.

is

the

The standard binary representation c b (n) is not a uniquely decodeable code for integers since there is no way of knowing when an integer has ended. For example, cb (5)cb (5) is identical to cb (45). It would be uniquely decodeable if we knew the standard binary length of each integer before it was received. Noticing that all positive integers have a standard binary representation that starts with a 1, we might define another representation: The headless binary representation of a positive integer n will be denoted by cB (n), e.g., cB (5) = 01, cB (45) = 01101 and cB (1) = λ (where λ denotes the null string). This representation would be uniquely decodeable if we knew the length l b (n) of the integer. So, how can we make a uniquely decodeable code for integers? Two strategies can be distinguished. 1. Self-delimiting codes. We first communicate somehow the length of the integer, lb (n), which is also a positive integer; then communicate the original integer n itself using cB (n). 2. Codes with ‘end of file’ characters. We code the integer into blocks of length b bits, and reserve one of the 2 b symbols to have the special meaning ‘end of file’. The coding of integers into blocks is arranged so that this reserved symbol is not needed for any other purpose. The simplest uniquely decodeable code for integers is the unary code, which can be viewed as a code with an end of file character. 132

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7 — Codes for Integers

133

Unary code. An integer n is encoded by sending a string of n−1 0s followed by a 1. n

cU (n)

1 2 3 4 5 .. .

1 01 001 0001 00001

45 000000000000000000000000000000000000000000001 The unary code has length lU (n) = n. The unary code is the optimal code for integers if the probability distribution over n is pU (n) = 2−n . Self-delimiting codes We can use the unary code to encode the length of the binary encoding of n and make a self-delimiting code: Code Cα . We send the unary code for lb (n), followed by the headless binary representation of n. cα (n) = cU [lb (n)]cB (n). (7.1) Table 7.1 shows the codes for some integers. The overlining indicates the division of each string into the parts c U [lb (n)] and cB (n). We might equivalently view cα (n) as consisting of a string of (lb (n) − 1) zeroes followed by the standard binary representation of n, c b (n).

n

cb (n) lb (n) cα (n)

1 2 3 4 5 6 .. . 45

1 10 11 100 101 110

1 2 2 3 3 3

1 010 011 00100 00101 00110

101101

6

00000101101

Table 7.1. Cα .

The codeword cα (n) has length lα (n) = 2lb (n) − 1.

The implicit probability distribution over n for the code C α is separable into the product of a probability distribution over the length l, P (l) = 2−l , and a uniform distribution over integers having that length,  −l+1 2 lb (n) = l P (n | l) = 0 otherwise.

(7.2)

(7.3)

Now, for the above code, the header that communicates the length always occupies the same number of bits as the standard binary representation of the integer (give or take one). If we are expecting to encounter large integers (large files) then this representation seems suboptimal, since it leads to all files occupying a size that is double their original uncoded size. Instead of using the unary code to encode the length l b (n), we could use Cα . Code Cβ . We send the length lb (n) using Cα , followed by the headless binary representation of n. cβ (n) = cα [lb (n)]cB (n). (7.4) Iterating this procedure, we can define a sequence of codes. Code Cγ . cγ (n) = cβ [lb (n)]cB (n).

(7.5)

cδ (n) = cγ [lb (n)]cB (n).

(7.6)

Code Cδ .

n

cβ (n)

cγ (n)

1 2 3 4 5 6 .. . 45

1 0100 0101 01100 01101 01110

1 01000 01001 010100 010101 010110

0011001101

0111001101

Table 7.2. Cβ and Cγ .

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134

7 — Codes for Integers

Codes with end-of-file symbols We can also make byte-based representations. (Let’s use the term byte flexibly here, to denote any fixed-length string of bits, not just a string of length 8 bits.) If we encode the number in some base, for example decimal, then we can represent each digit in a byte. In order to represent a digit from 0 to 9 in a byte we need four bits. Because 24 = 16, this leaves 6 extra four-bit symbols, {1010, 1011, 1100, 1101, 1110, 1111}, that correspond to no decimal digit. We can use these as end-of-file symbols to indicate the end of our positive integer. Clearly it is redundant to have more than one end-of-file symbol, so a more efficient code would encode the integer into base 15, and use just the sixteenth symbol, 1111, as the punctuation character. Generalizing this idea, we can make similar byte-based codes for integers in bases 3 and 7, and in any base of the form 2n − 1. These codes are almost complete. (Recall that a code is ‘complete’ if it satisfies the Kraft inequality with equality.) The codes’ remaining inefficiency is that they provide the ability to encode the integer zero and the empty string, neither of which was required. . Exercise 7.1.[2, p.136] Consider the implicit probability distribution over integers corresponding to the code with an end-of-file character. (a) If the code has eight-bit blocks (i.e., the integer is coded in base 255), what is the mean length in bits of the integer, under the implicit distribution? (b) If one wishes to encode binary files of expected size about one hundred kilobytes using a code with an end-of-file character, what is the optimal block size?

Encoding a tiny file To illustrate the codes we have discussed, we now use each code to encode a small file consisting of just 14 characters, Claude Shannon . • If we map the ASCII characters onto seven-bit symbols (e.g., in decimal, C = 67, l = 108, etc.), this 14 character file corresponds to the integer n = 167 987 786 364 950 891 085 602 469 870 (decimal). • The unary code for n consists of this many (less one) zeroes, followed by a one. If all the oceans were turned into ink, and if we wrote a hundred bits with every cubic millimeter, there might be enough ink to write cU (n). • The standard binary representation of n is this length-98 sequence of bits: cb (n) =

1000011110110011000011110101110010011001010100000 1010011110100011000011101110110111011011111101110.

. Exercise 7.2.[2 ] Write down or describe the following self-delimiting representations of the above number n: cα (n), cβ (n), cγ (n), cδ (n), c3 (n), c7 (n), and c15 (n). Which of these encodings is the shortest? [Answer: c 15 .]

n

c3 (n)

c7 (n)

1 2 3 .. .

01 11 10 11 01 00 11

001 111 010 111 011 111

45 01 10 00 00 11 110 011 111

Table 7.3. Two codes with end-of-file symbols, C3 and C7 . Spaces have been included to show the byte boundaries.

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7 — Codes for Integers

135

Comparing the codes One could answer the question ‘which of two codes is superior?’ by a sentence of the form ‘For n > k, code 1 is superior, for n < k, code 2 is superior’ but I contend that such an answer misses the point: any complete code corresponds to a prior for which it is optimal; you should not say that any other code is superior to it. Other codes are optimal for other priors. These implicit priors should be thought about so as to achieve the best code for one’s application. Notice that one cannot, for free, switch from one code to another, choosing whichever is shorter. If one were to do this, then it would be necessary to lengthen the message in some way that indicates which of the two codes is being used. If this is done by a single leading bit, it will be found that the resulting code is suboptimal because it fails the Kraft equality, as was discussed in exercise 5.33 (p.104). Another way to compare codes for integers is to consider a sequence of probability distributions, such as monotonic probability distributions over n ≥ 1, and rank the codes as to how well they encode any of these distributions. A code is called a ‘universal’ code if for any distribution in a given class, it encodes into an average length that is within some factor of the ideal average length. Let me say this again. We are meeting an alternative world view – rather than figuring out a good prior over integers, as advocated above, many theorists have studied the problem of creating codes that are reasonably good codes for any priors in a broad class. Here the class of priors conventionally considered is the set of priors that (a) assign a monotonically decreasing probability over integers and (b) have finite entropy. Several of the codes we have discussed above are universal. Another code which elegantly transcends the sequence of self-delimiting codes is Elias’s ‘universal code for integers’ (Elias, 1975), which effectively chooses from all the codes Cα , Cβ , . . . . It works by sending a sequence of messages each of which encodes the length of the next message, and indicates by a single bit whether or not that message is the final integer (in its standard binary representation). Because a length is a positive integer and all positive integers begin with ‘1’, all the leading 1s can be omitted. Write ‘0’ Loop { If blog nc = 0 halt Prepend cb (n) to the written string n:=blog nc }

The encoder of Cω is shown in algorithm 7.4. The encoding is generated from right to left. Table 7.5 shows the resulting codewords.

. Exercise 7.3.[2 ] Show that the Elias code is not actually the best code for a prior distribution that expects very large integers. (Do this by constructing another code and specifying how large n must be for your code to give a shorter length than Elias’s.)

Algorithm 7.4. Elias’s encoder for an integer n.

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136

7 — Codes for Integers

n

cω (n)

n

cω (n)

n

cω (n)

n

cω (n)

1 2 3 4 5 6 7 8

0 100 110 101000 101010 101100 101110 1110000

9 10 11 12 13 14 15 16

1110010 1110100 1110110 1111000 1111010 1111100 1111110 10100100000

31 32 45 63 64 127 128 255

10100111110 101011000000 101011011010 101011111110 1011010000000 1011011111110 10111100000000 10111111111110

256 365 511 512 719 1023 1024 1025

1110001000000000 1110001011011010 1110001111111110 11100110000000000 11100110110011110 11100111111111110 111010100000000000 111010100000000010

Solutions Solution to exercise 7.1 (p.134). The use of the end-of-file symbol in a code that represents the integer in some base q corresponds to a belief that there is a probability of (1/(q + 1)) that the current character is the last character of the number. Thus the prior to which this code is matched puts an exponential prior distribution over the length of the integer. (a) The expected number of characters is q +1 = 256, so the expected length of the integer is 256 × 8 ' 2000 bits. (b) We wish to find q such that q log q ' 800 000 bits. A value of q between 215 and 216 satisfies this constraint, so 16-bit blocks are roughly the optimal size, assuming there is one end-of-file character.

Table 7.5. Elias’s ‘universal’ code for integers. Examples from 1 to 1025.

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Part II

Noisy-Channel Coding

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8 Dependent Random Variables In the last three chapters on data compression we concentrated on random vectors x coming from an extremely simple probability distribution, namely the separable distribution in which each component x n is independent of the others. In this chapter, we consider joint ensembles in which the random variables are dependent. This material has two motivations. First, data from the real world have interesting correlations, so to do data compression well, we need to know how to work with models that include dependences. Second, a noisy channel with input x and output y defines a joint ensemble in which x and y are dependent – if they were independent, it would be impossible to communicate over the channel – so communication over noisy channels (the topic of chapters 9–11) is described in terms of the entropy of joint ensembles.

8.1 More about entropy This section gives definitions and exercises to do with entropy, carrying on from section 2.4. The joint entropy of X, Y is: H(X, Y ) =

X

P (x, y) log

xy∈AX AY

1 . P (x, y)

(8.1)

Entropy is additive for independent random variables: H(X, Y ) = H(X) + H(Y ) iff P (x, y) = P (x)P (y).

(8.2)

The conditional entropy of X given y = b k is the entropy of the probability distribution P (x | y = bk ). X 1 H(X | y = bk ) ≡ . (8.3) P (x | y = bk ) log P (x | y = bk ) x∈AX

The conditional entropy of X given Y is the average, over y, of the conditional entropy of X given y.   X X 1  H(X | Y ) ≡ P (y)  P (x | y) log P (x | y) y∈AY

=

X

x∈AX

P (x, y) log

xy∈AX AY

1 . P (x | y)

(8.4)

This measures the average uncertainty that remains about x when y is known. 138

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8.1: More about entropy

139

The marginal entropy of X is another name for the entropy of X, H(X), used to contrast it with the conditional entropies listed above. Chain rule for information content. From the product rule for probabilities, equation (2.6), we obtain: log

1 P (x, y)

= log

1 1 + log P (x) P (y | x)

(8.5)

so h(x, y) = h(x) + h(y | x).

(8.6)

In words, this says that the information content of x and y is the information content of x plus the information content of y given x. Chain rule for entropy. The joint entropy, marginal entropy are related by:

conditional entropy and

H(X, Y ) = H(X) + H(Y | X) = H(Y ) + H(X | Y ).

(8.7)

In words, this says that the uncertainty of X and Y is the uncertainty of X plus the uncertainty of Y given X. The mutual information between X and Y is I(X; Y ) ≡ H(X) − H(X | Y ),

(8.8)

and satisfies I(X; Y ) = I(Y ; X), and I(X; Y ) ≥ 0. It measures the average reduction in uncertainty about x that results from learning the value of y; or vice versa, the average amount of information that x conveys about y. The conditional mutual information between X and Y given z = c k is the mutual information between the random variables X and Y in the joint ensemble P (x, y | z = ck ), I(X; Y | z = ck ) = H(X | z = ck ) − H(X | Y, z = ck ).

(8.9)

The conditional mutual information between X and Y given Z is the average over z of the above conditional mutual information. I(X; Y | Z) = H(X | Z) − H(X | Y, Z).

(8.10)

No other ‘three-term entropies’ will be defined. For example, expressions such as I(X; Y ; Z) and I(X | Y ; Z) are illegal. But you may put conjunctions of arbitrary numbers of variables in each of the three spots in the expression I(X; Y | Z) – for example, I(A, B; C, D | E, F ) is fine: it measures how much information on average c and d convey about a and b, assuming e and f are known. Figure 8.1 shows how the total entropy H(X, Y ) of a joint ensemble can be broken down. This figure is important.



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140

8 — Dependent Random Variables Figure 8.1. The relationship between joint information, marginal entropy, conditional entropy and mutual entropy.

H(X, Y ) H(X) H(Y ) H(X | Y )

H(Y |X)

I(X; Y )

8.2 Exercises . Exercise 8.1.[1 ] Consider three independent random variables u, v, w with entropies Hu , Hv , Hw . Let X ≡ (U, V ) and Y ≡ (V, W ). What is H(X, Y )? What is H(X | Y )? What is I(X; Y )? . Exercise 8.2.[3, p.142] Referring to the definitions of conditional entropy (8.3– 8.4), confirm (with an example) that it is possible for H(X | y = b k ) to exceed H(X), but that the average, H(X | Y ), is less than H(X). So data are helpful – they do not increase uncertainty, on average. . Exercise 8.3.[2, p.143] Prove the chain rule for entropy, equation (8.7). [H(X, Y ) = H(X) + H(Y | X)]. Exercise 8.4.[2, p.143] Prove that the mutual information I(X; Y ) ≡ H(X) − H(X | Y ) satisfies I(X; Y ) = I(Y ; X) and I(X; Y ) ≥ 0. [Hint: see exercise 2.26 (p.37) and note that

I(X; Y ) = DKL (P (x, y)||P (x)P (y)).]

(8.11)

Exercise 8.5.[4 ] The ‘entropy distance’ between two random variables can be defined to be the difference between their joint entropy and their mutual information: DH (X, Y ) ≡ H(X, Y ) − I(X; Y ). (8.12) Prove that the entropy distance satisfies the axioms for a distance – DH (X, Y ) ≥ 0, DH (X, X) = 0, DH (X, Y ) = DH (Y, X), and DH (X, Z) ≤ DH (X, Y ) + DH (Y, Z). [Incidentally, we are unlikely to see D H (X, Y ) again but it is a good function on which to practise inequality-proving.] Exercise 8.6.[2, p.147] A joint ensemble XY has the following joint distribution. P (x, y)

y

1 2 3 4

x 1

2

3

4

1

1/8

1/16

1/32

1/32

2

1/16

1/8

1/32

1/32

3

1/16

1/16

1/16

1/16

4

1/4

0

0

0

1 2 3 4

What is the joint entropy H(X, Y )? What are the marginal entropies H(X) and H(Y )? For each value of y, what is the conditional entropy H(X | y)? What is the conditional entropy H(X | Y )? What is the conditional entropy of Y given X? What is the mutual information between X and Y ?

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8.3: Further exercises

141

Exercise 8.7.[2, p.143] Consider the ensemble XY Z in which A X = AY = AZ = {0, 1}, x and y are independent with P X = {p, 1 − p} and PY = {q, 1−q} and z = (x + y) mod 2. (8.13) (a) If q = 1/2, what is PZ ? What is I(Z; X)?

(b) For general p and q, what is PZ ? What is I(Z; X)? Notice that this ensemble is related to the binary symmetric channel, with x = input, y = noise, and z = output. Figure 8.2. A misleading representation of entropies (contrast with figure 8.1).

H(Y) H(X|Y) I(X;Y) H(Y|X)

H(X,Y)

H(X)

Three term entropies Exercise 8.8.[3, p.143] Many texts draw figure 8.1 in the form of a Venn diagram (figure 8.2). Discuss why this diagram is a misleading representation of entropies. Hint: consider the three-variable ensemble XY Z in which x ∈ {0, 1} and y ∈ {0, 1} are independent binary variables and z ∈ {0, 1} is defined to be z = x + y mod 2.

8.3 Further exercises The data-processing theorem The data processing theorem states that data processing can only destroy information. Exercise 8.9.[3, p.144] Prove this theorem by considering an ensemble W DR in which w is the state of the world, d is data gathered, and r is the processed data, so that these three variables form a Markov chain w → d → r,

(8.14)

that is, the probability P (w, d, r) can be written as P (w, d, r) = P (w)P (d | w)P (r | d).

(8.15)

Show that the average information that R conveys about W, I(W ; R), is less than or equal to the average information that D conveys about W , I(W ; D). This theorem is as much a caution about our definition of ‘information’ as it is a caution about data processing!

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142

8 — Dependent Random Variables

Inference and information measures Exercise 8.10.[2 ] The three cards. (a) One card is white on both faces; one is black on both faces; and one is white on one side and black on the other. The three cards are shuffled and their orientations randomized. One card is drawn and placed on the table. The upper face is black. What is the colour of its lower face? (Solve the inference problem.) (b) Does seeing the top face convey information about the colour of the bottom face? Discuss the information contents and entropies in this situation. Let the value of the upper face’s colour be u and the value of the lower face’s colour be l. Imagine that we draw a random card and learn both u and l. What is the entropy of u, H(U )? What is the entropy of l, H(L)? What is the mutual information between U and L, I(U ; L)?

Entropies of Markov processes . Exercise 8.11.[3 ] In the guessing game, we imagined predicting the next letter in a document starting from the beginning and working towards the end. Consider the task of predicting the reversed text, that is, predicting the letter that precedes those already known. Most people find this a harder task. Assuming that we model the language using an N -gram model (which says the probability of the next character depends only on the N − 1 preceding characters), is there any difference between the average information contents of the reversed language and the forward language?

8.4 Solutions Solution to exercise 8.2 (p.140). See exercise 8.6 (p.140) for an example where H(X | y) exceeds H(X) (set y = 3). We can prove the inequality H(X | Y ) ≤ H(X) by turning the expression into a relative entropy (using Bayes’ theorem) and invoking Gibbs’ inequality (exercise 2.26 (p.37)): H(X | Y ) ≡

X

y∈AY

P (y) 

X

=



X

x∈AX

 1  P (x | y) log P (x | y)

P (x, y) log

xy∈AX AY

=

X xy

=

X x

P (x)P (y | x) log P (x) log

1 P (x | y)

(8.16)

P (y) P (y | x)P (x)

(8.17)

X X 1 P (y) + . (8.18) P (x) P (y | x) log P (x) P (y | x) x y

The last expression is a sum of relative entropies between the distributions P (y | x) and P (y). So H(X | Y ) ≤ H(X) + 0, (8.19) with equality only if P (y | x) = P (y) for all x and y (that is, only if X and Y are independent).

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8.4: Solutions

143

Solution to exercise 8.3 (p.140). The chain rule for entropy follows from the decomposition of a joint probability: 1 (8.20) P (x, y) xy   X 1 1 + log (8.21) = P (x)P (y | x) log P (x) P (y | x) xy X X X 1 1 + P (x) P (y | x) log (8.22) = P (x) log P (x) P (y | x) x x y

H(X, Y ) =

X

P (x, y) log

= H(X) + H(Y | X).

Solution to exercise 8.4 (p.140).

(8.23)

Symmetry of mutual information:

I(X; Y ) = H(X) − H(X | Y ) X X 1 1 = P (x) log − P (x, y) log P (x) P (x | y) x xy =

X

X

(8.25)

P (x, y) log

P (x | y) P (x)

(8.26)

P (x, y) log

P (x, y) . P (x)P (y)

(8.27)

xy

=

(8.24)

xy

This expression is symmetric in x and y so I(X; Y ) = H(X) − H(X | Y ) = H(Y ) − H(Y | X).

(8.28)

We can prove that mutual information is positive in two ways. One is to continue from X P (x, y) (8.29) I(X; Y ) = P (x, y) log P (x)P (y) x,y which is a relative entropy and use Gibbs’ inequality (proved on p.44), which asserts that this relative entropy is ≥ 0, with equality only if P (x, y) = P (x)P (y), that is, if X and Y are independent. The other is to use Jensen’s inequality on −

X x,y

P (x, y) log

X P (x, y) P (x)P (y) ≥ − log P (x)P (y) = log 1 = 0. (8.30) P (x, y) P (x, y) x,y

Solution to exercise 8.7 (p.141).

z = x + y mod 2.

(a) If q = 1/2, PZ = {1/2, 1/2} and I(Z; X) = H(Z) − H(Z | X) = 1 − 1 = 0. (b) For general q and p, PZ = {pq+(1−p)(1−q), p(1−q)+q(1−p)}. The mutual information is I(Z; X) = H(Z)−H(Z | X) = H 2 (pq+(1−p)(1−q))−H2 (q).

Three term entropies Solution to exercise 8.8 (p.141). The depiction of entropies in terms of Venn diagrams is misleading for at least two reasons. First, one is used to thinking of Venn diagrams as depicting sets; but what are the ‘sets’ H(X) and H(Y ) depicted in figure 8.2, and what are the objects that are members of those sets? I think this diagram encourages the novice student to make inappropriate analogies. For example, some students imagine

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144 I(X;Y)

H(X|Y,Z) H(X)

8 — Dependent Random Variables

                                                                                                                      

A H(Y) H(Y|X,Z)

H(Z|X)

I(X;Y|Z)

                                                                                           

H(Z|Y)

H(X,Y|Z)

H(Z|X,Y)

H(Z)

that the random outcome (x, y) might correspond to a point in the diagram, and thus confuse entropies with probabilities. Secondly, the depiction in terms of Venn diagrams encourages one to believe that all the areas correspond to positive quantities. In the special case of two random variables it is indeed true that H(X | Y ), I(X; Y ) and H(Y | X) are positive quantities. But as soon as we progress to three-variable ensembles, we obtain a diagram with positive-looking areas that may actually correspond to negative quantities. Figure 8.3 correctly shows relationships such as H(X) + H(Z | X) + H(Y | X, Z) = H(X, Y, Z).

(8.31)

But it gives the misleading impression that the conditional mutual information I(X; Y | Z) is less than the mutual information I(X; Y ). In fact the area labelled A can correspond to a negative quantity. Consider the joint ensemble (X, Y, Z) in which x ∈ {0, 1} and y ∈ {0, 1} are independent binary variables and z ∈ {0, 1} is defined to be z = x + y mod 2. Then clearly H(X) = H(Y ) = 1 bit. Also H(Z) = 1 bit. And H(Y | X) = H(Y ) = 1 since the two variables are independent. So the mutual information between X and Y is zero. I(X; Y ) = 0. However, if z is observed, X and Y become dependent — knowing x, given z, tells you what y is: y = z − x mod 2. So I(X; Y | Z) = 1 bit. Thus the area labelled A must correspond to −1 bits for the figure to give the correct answers. The above example is not at all a capricious or exceptional illustration. The binary symmetric channel with input X, noise Y , and output Z is a situation in which I(X; Y ) = 0 (input and noise are independent) but I(X; Y | Z) > 0 (once you see the output, the unknown input and the unknown noise are intimately related!). The Venn diagram representation is therefore valid only if one is aware that positive areas may represent negative quantities. With this proviso kept in mind, the interpretation of entropies in terms of sets can be helpful (Yeung, 1991). Solution to exercise 8.9 (p.141). For any joint ensemble XY Z, the following chain rule for mutual information holds. I(X; Y, Z) = I(X; Y ) + I(X; Z | Y ).

(8.32)

Now, in the case w → d → r, w and r are independent given d, so I(W ; R | D) = 0. Using the chain rule twice, we have: I(W ; D, R) = I(W ; D)

(8.33)

I(W ; D, R) = I(W ; R) + I(W ; D | R),

(8.34)

I(W ; R) − I(W ; D) ≤ 0.

(8.35)

and so

Figure 8.3. A misleading representation of entropies, continued.

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About Chapter 9 Before reading Chapter 9, you should have read Chapter 1 and worked on exercise 2.26 (p.37), and exercises 8.2–8.7 (pp.140–141).

145

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9 Communication over a Noisy Channel 9.1 The big picture

Source 6

?

Source coding

Compressor

Decompressor 6

?

Channel coding

Encoder

Decoder Noisy channel

-

6

In Chapters 4–6, we discussed source coding with block codes, symbol codes and stream codes. We implicitly assumed that the channel from the compressor to the decompressor was noise-free. Real channels are noisy. We will now spend two chapters on the subject of noisy-channel coding – the fundamental possibilities and limitations of error-free communication through a noisy channel. The aim of channel coding is to make the noisy channel behave like a noiseless channel. We will assume that the data to be transmitted has been through a good compressor, so the bit stream has no obvious redundancy. The channel code, which makes the transmission, will put back redundancy of a special sort, designed to make the noisy received signal decodeable. Suppose we transmit 1000 bits per second with p 0 = p1 = 1/2 over a noisy channel that flips bits with probability f = 0.1. What is the rate of transmission of information? We might guess that the rate is 900 bits per second by subtracting the expected number of errors per second. But this is not correct, because the recipient does not know where the errors occurred. Consider the case where the noise is so great that the received symbols are independent of the transmitted symbols. This corresponds to a noise level of f = 0.5, since half of the received symbols are correct due to chance alone. But when f = 0.5, no information is transmitted at all. Given what we have learnt about entropy, it seems reasonable that a measure of the information transmitted is given by the mutual information between the source and the received signal, that is, the entropy of the source minus the conditional entropy of the source given the received signal. We will now review the definition of conditional entropy and mutual information. Then we will examine whether it is possible to use such a noisy channel to communicate reliably. We will show that for any channel Q there is a non-zero rate, the capacity C(Q), up to which information can be sent 146

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9.2: Review of probability and information

147

with arbitrarily small probability of error.

9.2 Review of probability and information As an example, we take the joint distribution XY from exercise 8.6 (p.140). The marginal distributions P (x) and P (y) are shown in the margins. P (x, y)

x

P (y)

1

2

3

4

1 2 3 4

1/8 1/16

1/16

1/32

1/32

1/4

1/8

1/32

1/32

1/4

1/16

1/16

1/16

1/16

1/4

1/4

0

0

0

1/4

P (x)

1/2

1/4

1/8

1/8

y

The joint entropy is H(X, Y ) = 27/8 bits. The marginal entropies are H(X) = 7/4 bits and H(Y ) = 2 bits. We can compute the conditional distribution of x for each value of y, and the entropy of each of those conditional distributions: P (x | y)

y

1 2 3 4

H(X | y)/bits

x 1

2

3

4

1/2

1/4

1/8

1/8

7/4

1/4

1/2

1/8

1/8

7/4

1/4

1/4

1/4

1/4

1

0

0

0

2 0

H(X | Y ) = 11/8 Note that whereas H(X | y = 4) = 0 is less than H(X), H(X | y = 3) is greater than H(X). So in some cases, learning y can increase our uncertainty about x. Note also that although P (x | y = 2) is a different distribution from P (x), the conditional entropy H(X | y = 2) is equal to H(X). So learning that y is 2 changes our knowledge about x but does not reduce the uncertainty of x, as measured by the entropy. On average though, learning y does convey information about x, since H(X | Y ) < H(X). One may also evaluate H(Y |X) = 13/8 bits. The mutual information is I(X; Y ) = H(X) − H(X | Y ) = 3/8 bits.

9.3 Noisy channels A discrete memoryless channel Q is characterized by an input alphabet AX , an output alphabet AY , and a set of conditional probability distributions P (y | x), one for each x ∈ AX . These transition probabilities may be written in a matrix Qj|i = P (y = bj | x = ai ).

(9.1)

I usually orient this matrix with the output variable j indexing the rows and the input variable i indexing the columns, so that each column of Q is a probability vector. With this convention, we can obtain the probability of the output, pY , from a probability distribution over the input, pX , by right-multiplication: pY = QpX . (9.2)

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148

9 — Communication over a Noisy Channel

Some useful model channels are: Binary symmetric channel. AX = {0, 1}. AY = {0, 1}. x

0 -0 @ 

R @ 1 -1

0 1

P (y = 0 | x = 0) = 1 − f ; P (y = 0 | x = 1) = f ; P (y = 1 | x = 0) = f ; P (y = 1 | x = 1) = 1 − f.

y

0 1

Binary erasure channel. AX = {0, 1}. AY = {0, ?, 1}. 0 -0 R @? y x @  1 -1

0 1

P (y = 0 | x = 0) = 1 − f ; P (y = 0 | x = 1) = 0; P (y = ? | x = 0) = f ; P (y = ? | x = 1) = f ; P (y = 1 | x = 0) = 0; P (y = 1 | x = 1) = 1 − f.

0 ? 1

Noisy typewriter. AX = AY = the 27 letters {A, B, . . . , Z, -}. The letters are arranged in a circle, and when the typist attempts to type B, what comes out is either A, B or C, with probability 1/3 each; when the input is C, the output is B, C or D; and so forth, with the final letter ‘-’ adjacent to the first letter A. A A PP 1 q P 1  B P B C Pq  P C C C P 1  Pq P  D 1  D PC  Pq P  E 1  PC E P q PF 1 PC F  P q PG C 1  G P P C q P PP 1 H H P  .. Cq . C P P1 q P C Y Y P P 1   1 q PC Z Z  P Pq -WC   P -

ABCDEFGHIJKLMNOPQRSTUVWXYZA B C D E F G H I J K L M N O P Q R S T U V W X Y Z -

.. . P (y = F | x = G) = 1/3; P (y = G | x = G) = 1/3; P (y = H | x = G) = 1/3; .. .

Z channel. AX = {0, 1}. AY = {0, 1}. x

0 -0 

1 -1

0 1

P (y = 0 | x = 0) = 1; P (y = 0 | x = 1) = f ; P (y = 1 | x = 0) = 0; P (y = 1 | x = 1) = 1 − f.

y

9.4 Inferring the input given the output If we assume that the input x to a channel comes from an ensemble X, then we obtain a joint ensemble XY in which the random variables x and y have the joint distribution: P (x, y) = P (y | x)P (x). (9.3)

Now if we receive a particular symbol y, what was the input symbol x? We typically won’t know for certain. We can write down the posterior distribution of the input using Bayes’ theorem: P (x | y) =

P (y | x)P (x) P (y | x)P (x) =P . 0 0 P (y) x0 P (y | x )P (x )

(9.4)

Example 9.1. Consider a binary symmetric channel with probability of error f = 0.15. Let the input ensemble be PX : {p0 = 0.9, p1 = 0.1}. Assume we observe y = 1. P (x = 1 | y = 1)

= = =

P (y = 1 | x = 1)P (x = 1) P 0 0 x0 P (y | x )P (x ) 0.85 × 0.1 0.85 × 0.1 + 0.15 × 0.9 0.085 = 0.39. 0.22

(9.5)

0 1

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9.5: Information conveyed by a channel

149

Thus ‘x = 1’ is still less probable than ‘x = 0’, although it is not as improbable as it was before. Exercise 9.2.[1, p.157] Now assume we observe y = 0. Compute the probability of x = 1 given y = 0. Example 9.3. Consider a Z channel with probability of error f = 0.15. Let the input ensemble be PX : {p0 = 0.9, p1 = 0.1}. Assume we observe y = 1. P (x = 1 | y = 1) = =

0.85 × 0.1 0.85 × 0.1 + 0 × 0.9 0.085 = 1.0. 0.085

(9.6)

So given the output y = 1 we become certain of the input. Exercise 9.4.[1, p.157] Alternatively, assume we observe y = 0. P (x = 1 | y = 0).

Compute

9.5 Information conveyed by a channel We now consider how much information can be communicated through a channel. In operational terms, we are interested in finding ways of using the channel such that all the bits that are communicated are recovered with negligible probability of error. In mathematical terms, assuming a particular input ensemble X, we can measure how much information the output conveys about the input by the mutual information: I(X; Y ) ≡ H(X) − H(X | Y ) = H(Y ) − H(Y |X).

(9.7)

Our aim is to establish the connection between these two ideas. Let us evaluate I(X; Y ) for some of the channels above.

Hint for computing mutual information We will tend to think of I(X; Y ) as H(X) − H(X | Y ), i.e., how much the uncertainty of the input X is reduced when we look at the output Y . But for computational purposes it is often handy to evaluate H(Y )−H(Y |X) instead. H(X, Y ) H(X) H(Y ) H(X | Y )

I(X; Y )

H(Y |X)

Example 9.5. Consider the binary symmetric channel again, with f = 0.15 and PX : {p0 = 0.9, p1 = 0.1}. We already evaluated the marginal probabilities P (y) implicitly above: P (y = 0) = 0.78; P (y = 1) = 0.22. The mutual information is: I(X; Y ) = H(Y ) − H(Y |X).

Figure 9.1. The relationship between joint information, marginal entropy, conditional entropy and mutual entropy. This figure is important, so I’m showing it twice.

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150

9 — Communication over a Noisy Channel What is H(Y |X)? It is defined to be the weighted sum over x of H(Y | x); but H(Y | x) is the same for each value of x: H(Y | x = 0) is H 2 (0.15), and H(Y | x = 1) is H2 (0.15). So I(X; Y ) = H(Y ) − H(Y |X)

= H2 (0.22) − H2 (0.15)

= 0.76 − 0.61 = 0.15 bits.

(9.8)

This may be contrasted with the entropy of the source H(X) = H2 (0.1) = 0.47 bits. Note: here we have used the binary entropy function H 2 (p) ≡ H(p, 1− 1 . p) = p log p1 + (1 − p) log (1−p) Example 9.6. And now the Z channel, with P X as above. P (y = 1) = 0.085. I(X; Y ) = H(Y ) − H(Y |X)

= H2 (0.085) − [0.9H2 (0) + 0.1H2 (0.15)] = 0.42 − (0.1 × 0.61) = 0.36 bits.

(9.9)

The entropy of the source, as above, is H(X) = 0.47 bits. Notice that the mutual information I(X; Y ) for the Z channel is bigger than the mutual information for the binary symmetric channel with the same f . The Z channel is a more reliable channel. Exercise 9.7.[1, p.157] Compute the mutual information between X and Y for the binary symmetric channel with f = 0.15 when the input distribution is PX = {p0 = 0.5, p1 = 0.5}. Exercise 9.8.[2, p.157] Compute the mutual information between X and Y for the Z channel with f = 0.15 when the input distribution is P X : {p0 = 0.5, p1 = 0.5}.

Maximizing the mutual information We have observed in the above examples that the mutual information between the input and the output depends on the chosen input ensemble. Let us assume that we wish to maximize the mutual information conveyed by the channel by choosing the best possible input ensemble. We define the capacity of the channel to be its maximum mutual information. The capacity of a channel Q is: C(Q) = max I(X; Y ). PX

(9.10)

The distribution PX that achieves the maximum is called the optimal ∗ . [There may be multiple optimal input distribution, denoted by PX input distributions achieving the same value of I(X; Y ).] In Chapter 10 we will show that the capacity does indeed measure the maximum amount of error-free information that can be transmitted over the channel per unit time. Example 9.9. Consider the binary symmetric channel with f = 0.15. Above, we considered PX = {p0 = 0.9, p1 = 0.1}, and found I(X; Y ) = 0.15 bits.

Throughout this book, log means log2 .

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9.6: The noisy-channel coding theorem

151

How much better can we do? By symmetry, the optimal input distribution is {0.5, 0.5} and the capacity is C(QBSC ) = H2 (0.5) − H2 (0.15) = 1.0 − 0.61 = 0.39 bits.

(9.11)

We’ll justify the symmetry argument later. If there’s any doubt about the symmetry argument, we can always resort to explicit maximization of the mutual information I(X; Y ),

I(X; Y ) 0.4 0.3 0.2 0.1

I(X; Y ) = H2 ((1−f )p1 + (1−p1 )f ) − H2 (f )

(figure 9.2).

(9.12) 0 0

Example 9.10. The noisy typewriter. The optimal input distribution is a uniform distribution over x, and gives C = log 2 9 bits. Example 9.11. Consider the Z channel with f = 0.15. Identifying the optimal input distribution is not so straightforward. We evaluate I(X; Y ) explicitly for PX = {p0 , p1 }. First, we need to compute P (y). The probability of y = 1 is easiest to write down: P (y = 1) = p1 (1 − f ).

0.5

0.75

1

p1 Figure 9.2. The mutual information I(X; Y ) for a binary symmetric channel with f = 0.15 as a function of the input distribution.

(9.13)

Then the mutual information is:

I(X; Y )

I(X; Y ) = H(Y ) − H(Y |X)

0.7

= H2 (p1 (1 − f )) − (p0 H2 (0) + p1 H2 (f )) = H2 (p1 (1 − f )) − p1 H2 (f ).

0.6

(9.14)

This is a non-trivial function of p1 , shown in figure 9.3. It is maximized for f = 0.15 by p∗1 = 0.445. We find C(QZ ) = 0.685. Notice the optimal input distribution is not {0.5, 0.5}. We can communicate slightly more information by using input symbol 0 more frequently than 1. Exercise 9.12.[1, p.158] What is the capacity of the binary symmetric channel for general f ? Exercise 9.13.[2, p.158] Show that the capacity of the binary erasure channel with f = 0.15 is CBEC = 0.85. What is its capacity for general f ? Comment.

9.6 The noisy-channel coding theorem It seems plausible that the ‘capacity’ we have defined may be a measure of information conveyed by a channel; what is not obvious, and what we will prove in the next chapter, is that the capacity indeed measures the rate at which blocks of data can be communicated over the channel with arbitrarily small probability of error. We make the following definitions. An (N, K) block code for a channel Q is a list of S = 2 K codewords {x(1) , x(2) , . . . , x(2

0.25

K)

},

x(s) ∈ AN X,

each of length N . Using this code we can encode a signal s ∈ {1, 2, 3, . . . , 2K } as x(s) . [The number of codewords S is an integer, but the number of bits specified by choosing a codeword, K ≡ log 2 S, is not necessarily an integer.]

0.5 0.4 0.3 0.2 0.1 0 0

0.25

0.5

0.75

1

p1 Figure 9.3. The mutual information I(X; Y ) for a Z channel with f = 0.15 as a function of the input distribution.

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152

9 — Communication over a Noisy Channel The rate of the code is R = K/N bits per channel use. [We will use this definition of the rate for any channel, not only channels with binary inputs; note however that it is sometimes conventional to define the rate of a code for a channel with q input symbols to be K/(N log q).]

A decoder for an (N, K) block code is a mapping from the set of length-N strings of channel outputs, AN ˆ ∈ {0, 1, 2, . . . , 2 K }. Y , to a codeword label s The extra symbol sˆ = 0 can be used to indicate a ‘failure’.

The probability of block error of a code and decoder, for a given channel, and for a given probability distribution over the encoded signal P (s in ), is: X P (sin )P (sout 6= sin | sin ). (9.15) pB = sin

The maximal probability of block error is pBM = max P (sout 6= sin | sin ). sin

(9.16)

The optimal decoder for a channel code is the one that minimizes the probability of block error. It decodes an output y as the input s that has maximum posterior probability P (s | y). P (s | y) = P

P (y | s)P (s) 0 0 s0 P (y | s )P (s )

sˆoptimal = argmax P (s | y).

(9.17) (9.18)

A uniform prior distribution on s is usually assumed, in which case the optimal decoder is also the maximum likelihood decoder, i.e., the decoder that maps an output y to the input s that has maximum likelihood P (y | s). The probability of bit error pb is defined assuming that the codeword number s is represented by a binary vector s of length K bits; it is the average probability that a bit of sout is not equal to the corresponding bit of sin (averaging over all K bits). Shannon’s noisy-channel coding theorem (part one). Associated with each discrete memoryless channel, there is a non-negative number C (called the channel capacity) with the following property. For any  > 0 and R < C, for large enough N , there exists a block code of length N and rate ≥ R and a decoding algorithm, such that the maximal probability of block error is < .

Confirmation of the theorem for the noisy typewriter channel In the case of the noisy typewriter, we can easily confirm the theorem, because we can create a completely error-free communication strategy using a block code of length N = 1: we use only the letters B, E, H, . . . , Z, i.e., every third letter. These letters form a non-confusable subset of the input alphabet (see figure 9.5). Any output can be uniquely decoded. The number of inputs in the non-confusable subset is 9, so the error-free information rate of this system is log 2 9 bits, which is equal to the capacity C, which we evaluated in example 9.10 (p.151).

pBM 6

achievable -

C

R

Figure 9.4. Portion of the R, pBM plane asserted to be achievable by the first part of Shannon’s noisy channel coding theorem.

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9.7: Intuitive preview of proof

153 ABCDEFGHIJKLMNOPQRSTUVWXYZ-

A 1 PP- B B q PC 1  D PP- E E q P

A B C D E F G H I J K L M N O P Q R S T U V W X Y Z -

F

G 1 PP - H H q P . I

..

00 10 01 11

0000 1000 0100 1100 0010 1010 0110 1110 0001 1001 0101 1101 0011 1011 0111 1111

1 Y  - Z ZP Pq P-

0 1 0 1

N =1

00 10 01 11

0000 1000 0100 1100 0010 1010 0110 1110 0001 1001 0101 1101 0011 1011 0111 1111

N =2

Figure 9.5. A non-confusable subset of inputs for the noisy typewriter.

Figure 9.6. Extended channels obtained from a binary symmetric channel with transition probability 0.15.

N =4

How does this translate into the terms of the theorem? The following table explains. The theorem

How it applies to the noisy typewriter

Associated with each discrete memoryless channel, there is a non-negative number C. For any  > 0 and R < C, for large enough N , there exists a block code of length N and rate ≥ R

The capacity C is log 2 9.

No matter what  and R are, we set the blocklength N to 1. The block code is {B, E, . . . , Z}. The value of K is given by 2K = 9, so K = log 2 9, and this code has rate log 2 9, which is greater than the requested value of R.

and a decoding algorithm,

The decoding algorithm maps the received letter to the nearest letter in the code;

such that the maximal probability of block error is < .

the maximal probability of block error is zero, which is less than the given .

9.7 Intuitive preview of proof Extended channels To prove the theorem for any given channel, we consider the extended channel corresponding to N uses of the channel. The extended channel has |A X |N possible inputs x and |AY |N possible outputs. Extended channels obtained from a binary symmetric channel and from a Z channel are shown in figures 9.6 and 9.7, with N = 2 and N = 4.

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9 — Communication over a Noisy Channel

00 10 01 11

0000 1000 0100 1100 0010 1010 0110 1110 0001 1001 0101 1101 0011 1011 0111 1111

154

0 1 0 1

N =1 AN Y

00 10 01 11

0000 1000 0100 1100 0010 1010 0110 1110 0001 1001 0101 1101 0011 1011 0111 1111

N =2

Typical y ' $                                                                                            6        & %

Typical y for a given typical x (a)

Figure 9.7. Extended channels obtained from a Z channel with transition probability 0.15. Each column corresponds to an input, and each row is a different output.

N =4 AN Y

Typical y ' $         

           

  % & (b)

Exercise 9.14.[2, p.159] Find the transition probability matrices Q for the extended channel, with N = 2, derived from the binary erasure channel having erasure probability 0.15. By selecting two columns of this transition probability matrix, we can define a rate-1/2 code for this channel with blocklength N = 2. What is the best choice of two columns? What is the decoding algorithm? To prove the noisy-channel coding theorem, we make use of large blocklengths N . The intuitive idea is that, if N is large, an extended channel looks a lot like the noisy typewriter. Any particular input x is very likely to produce an output in a small subspace of the output alphabet – the typical output set, given that input. So we can find a non-confusable subset of the inputs that produce essentially disjoint output sequences. For a given N , let us consider a way of generating such a non-confusable subset of the inputs, and count up how many distinct inputs it contains. Imagine making an input sequence x for the extended channel by drawing it from an ensemble X N , where X is an arbitrary ensemble over the input alphabet. Recall the source coding theorem of Chapter 4, and consider the number of probable output sequences y. The total number of typical output sequences y is 2N H(Y ) , all having similar probability. For any particular typical input sequence x, there are about 2N H(Y |X) probable sequences. Some of these subsets of AN Y are depicted by circles in figure 9.8a. We now imagine restricting ourselves to a subset of the typical inputs x such that the corresponding typical output sets do not overlap, as shown in figure 9.8b. We can then bound the number of non-confusable inputs by dividing the size of the typical y set, 2 N H(Y ) , by the size of each typical-y-

Figure 9.8. (a) Some typical outputs in AN Y corresponding to typical inputs x. (b) A subset of the typical sets shown in (a) that do not overlap each other. This picture can be compared with the solution to the noisy typewriter in figure 9.5.

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9.8: Further exercises

155

given-typical-x set, 2N H(Y |X) . So the number of non-confusable inputs, if they are selected from the set of typical inputs x ∼ X N , is ≤ 2N H(Y )−N H(Y |X) = 2N I(X;Y ) . The maximum value of this bound is achieved if X is the ensemble that maximizes I(X; Y ), in which case the number of non-confusable inputs is ≤ 2N C . Thus asymptotically up to C bits per cycle, and no more, can be communicated with vanishing error probability. 2 This sketch has not rigorously proved that reliable communication really is possible – that’s our task for the next chapter.

9.8 Further exercises Exercise 9.15.[3, p.159] Refer back to the computation of the capacity of the Z channel with f = 0.15. (a) Why is p∗1 less than 0.5? One could argue that it is good to favour the 0 input, since it is transmitted without error – and also argue that it is good to favour the 1 input, since it often gives rise to the highly prized 1 output, which allows certain identification of the input! Try to make a convincing argument. (b) In the case of general f , show that the optimal input distribution is 1/(1 − f ) p∗1 = . (9.19) 1 + 2(H2 (f )/(1−f )) (c) What happens to p∗1 if the noise level f is very close to 1? Exercise 9.16.[2, p.159] Sketch graphs of the capacity of the Z channel, the binary symmetric channel and the binary erasure channel as a function of f . . Exercise 9.17.[2 ] What is the capacity of the five-input, ten-output channel whose transition probability matrix is                

0.25 0.25 0.25 0.25 0 0 0 0 0 0

0 0 0.25 0.25 0.25 0.25 0 0 0 0

0 0 0 0 0.25 0.25 0.25 0.25 0 0

0 0 0 0 0 0 0.25 0.25 0.25 0.25

0.25 0.25 0 0 0 0 0 0 0.25 0.25

               

01234 0 1 2 3 4 5 6 7 8 9

?

(9.20)

Exercise 9.18.[2, p.159] Consider a Gaussian channel with binary input x ∈ {−1, +1} and real output alphabet AY , with transition probability density (y−xα)2 1 − 2σ 2 , √ Q(y | x, α, σ) = (9.21) e 2πσ 2 where α is the signal amplitude. (a) Compute the posterior probability of x given y, assuming that the two inputs are equiprobable. Put your answer in the form P (x = 1 | y, α, σ) =

1 . 1 + e−a(y)

(9.22)

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156

9 — Communication over a Noisy Channel Sketch the value of P (x = 1 | y, α, σ) as a function of y.

(b) Assume that a single bit is to be transmitted. What is the optimal decoder, and what is its probability of error? Express your answer in terms of the signal-to-noise ratio α 2 /σ 2 and the error function (the cumulative probability function of the Gaussian distribution), Z z z2 1 √ e− 2 dz. (9.23) Φ(z) ≡ 2π −∞ [Note that this definition of the error function Φ(z) may not correspond to other people’s.]

Pattern recognition as a noisy channel We may think of many pattern recognition problems in terms of communication channels. Consider the case of recognizing handwritten digits (such as postcodes on envelopes). The author of the digit wishes to communicate a message from the set AX = {0, 1, 2, 3, . . . , 9}; this selected message is the input to the channel. What comes out of the channel is a pattern of ink on paper. If the ink pattern is represented using 256 binary pixels, the channel Q has as its output a random variable y ∈ A Y = {0, 1}256 . An example of an element from this alphabet is shown in the margin. Exercise 9.19.[2 ] Estimate how many patterns in AY are recognizable as the character ‘2’. [The aim of this problem is to try to demonstrate the existence of as many patterns as possible that are recognizable as 2s.] Discuss how one might model the channel P (y | x = 2). Estimate the entropy of the probability distribution P (y | x = 2). One strategy for doing pattern recognition is to create a model for P (y | x) for each value of the input x = {0, 1, 2, 3, . . . , 9}, then use Bayes’ theorem to infer x given y. P (x | y) = P

P (y | x)P (x) . 0 0 x0 P (y | x )P (x )

(9.24)

This strategy is known as full probabilistic modelling or generative modelling. This is essentially how current speech recognition systems work. In addition to the channel model, P (y | x), one uses a prior probability distribution P (x), which in the case of both character recognition and speech recognition is a language model that specifies the probability of the next character/word given the context and the known grammar and statistics of the language.

Random coding Exercise 9.20.[2, p.160] Given twenty-four people in a room, what is the probability that there are at least two people present who have the same birthday (i.e., day and month of birth)? What is the expected number of pairs of people with the same birthday? Which of these two questions is easiest to solve? Which answer gives most insight? You may find it helpful to solve these problems and those that follow using notation such as A = number of days in year = 365 and S = number of people = 24. . Exercise 9.21.[2 ] The birthday problem may be related to a coding scheme. Assume we wish to convey a message to an outsider identifying one of

Figure 9.9. Some more 2s.

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9.9: Solutions

157

the twenty-four people. We could simply communicate a number s from AS = {1, 2, . . . , 24}, having agreed a mapping of people onto numbers; alternatively, we could convey a number from A X = {1, 2, . . . , 365}, identifying the day of the year that is the selected person’s birthday (with apologies to leapyearians). [The receiver is assumed to know all the people’s birthdays.] What, roughly, is the probability of error of this communication scheme, assuming it is used for a single transmission? What is the capacity of the communication channel, and what is the rate of communication attempted by this scheme? . Exercise 9.22.[2 ] Now imagine that there are K rooms in a building, each containing q people. (You might think of K = 2 and q = 24 as an example.) The aim is to communicate a selection of one person from each room by transmitting an ordered list of K days (from A X ). Compare the probability of error of the following two schemes. (a) As before, where each room transmits the birthday of the selected person. (b) To each K-tuple of people, one drawn from each room, an ordered K-tuple of randomly selected days from A X is assigned (this Ktuple has nothing to do with their birthdays). This enormous list of S = q K strings is known to the receiver. When the building has selected a particular person from each room, the ordered string of days corresponding to that K-tuple of people is transmitted. What is the probability of error when q = 364 and K = 1? What is the probability of error when q = 364 and K is large, e.g. K = 6000?

9.9 Solutions Solution to exercise 9.2 (p.149).

If we assume we observe y = 0,

P (x = 1 | y = 0) = = = Solution to exercise 9.4 (p.149).

(9.25) (9.26) (9.27)

If we observe y = 0,

P (x = 1 | y = 0) = = Solution to exercise 9.7 (p.150). mutual information is:

P (y = 0 | x = 1)P (x = 1) P 0 0 x0 P (y | x )P (x ) 0.15 × 0.1 0.15 × 0.1 + 0.85 × 0.9 0.015 = 0.019. 0.78

0.15 × 0.1 0.15 × 0.1 + 1.0 × 0.9 0.015 = 0.016. 0.915

(9.28) (9.29)

The probability that y = 1 is 0.5, so the

I(X; Y ) = H(Y ) − H(Y | X)

= H2 (0.5) − H2 (0.15)

= 1 − 0.61 = 0.39 bits.

(9.30) (9.31) (9.32)

Solution to exercise 9.8 (p.150). We again compute the mutual information using I(X; Y ) = H(Y ) − H(Y | X). The probability that y = 0 is 0.575, and

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158

9 — Communication over a Noisy Channel

P H(Y | X) = x P (x)H(Y | x) = P (x = 1)H(Y | x = 1) + P (x = 0)H(Y | x = 0) so the mutual information is: I(X; Y ) = H(Y ) − H(Y | X)

= H2 (0.575) − [0.5 × H2 (0.15) + 0.5 × 0] = 0.98 − 0.30 = 0.679 bits.

(9.33) (9.34) (9.35)

Solution to exercise 9.12 (p.151). By symmetry, the optimal input distribution is {0.5, 0.5}. Then the capacity is C = I(X; Y ) = H(Y ) − H(Y | X) = H2 (0.5) − H2 (f ) = 1 − H2 (f ).

(9.36) (9.37) (9.38)

Would you like to find the optimal input distribution without invoking symmetry? We can do this by computing the mutual information in the general case where the input ensemble is {p0 , p1 }: I(X; Y ) = H(Y ) − H(Y | X)

= H2 (p0 f + p1 (1 − f )) − H2 (f ).

(9.39) (9.40)

The only p-dependence is in the first term H 2 (p0 f + p1 (1 − f )), which is maximized by setting the argument to 0.5. This value is given by setting p0 = 1/2. Solution to exercise 9.13 (p.151). Answer 1. By symmetry, the optimal input distribution is {0.5, 0.5}. The capacity is most easily evaluated by writing the mutual information as I(X; Y ) = H(X) − H(X | Y ). The conditional entropy P H(X | Y ) is y P (y)H(X | y); when y is known, x is uncertain only if y = ?, which occurs with probability f /2 + f /2, so the conditional entropy H(X | Y ) is f H2 (0.5). C = I(X; Y ) = H(X) − H(X | Y )

= H2 (0.5) − f H2 (0.5)

= 1 − f.

(9.41) (9.42) (9.43)

The binary erasure channel fails a fraction f of the time. Its capacity is precisely 1 − f , which is the fraction of the time that the channel is reliable. This result seems very reasonable, but it is far from obvious how to encode information so as to communicate reliably over this channel. Answer 2. Alternatively, without invoking the symmetry assumed above, we can start from the input ensemble {p 0 , p1 }. The probability that y = ? is p0 f + p1 f = f , and when we receive y = ?, the posterior probability of x is the same as the prior probability, so: I(X; Y ) = H(X) − H(X | Y ) = H2 (p1 ) − f H2 (p1 ) = (1 − f )H2 (p1 ).

(9.44) (9.45) (9.46)

This mutual information achieves its maximum value of (1−f ) when p 1 = 1/2.

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9.9: Solutions

159

0 1

Q

0 ? 1

(a)

N =1

00 ?0 10 0? ?? 1? 01 ?1 11

(b)

00 10 01 11

x(1) x(2)

00 10 01 11

00 10 01 11

x(1) x(2)

00 ?0 10 0? ?? 1? 01 ?1 11

(c)

00 ?0 10 0? ?? 1? 01 ?1 11

-

m ˆ =1 m ˆ =1

-

m ˆ =1 m ˆ =0 m ˆ =2

-

m ˆ =2 m ˆ =2

Figure 9.10. (a) The extended channel (N = 2) obtained from a binary erasure channel with erasure probability 0.15. (b) A block code consisting of the two codewords 00 and 11. (c) The optimal decoder for this code.

N =2

Solution to exercise 9.14 (p.153). The extended channel is shown in figure 9.10. The best code for this channel with N = 2 is obtained by choosing two columns that have minimal overlap, for example, columns 00 and 11. The decoding algorithm returns ‘00’ if the extended channel output is among the top four and ‘11’ if it’s among the bottom four, and gives up if the output is ‘??’. Solution to exercise 9.15 (p.155). In example 9.11 (p.151) we showed that the mutual information between input and output of the Z channel is I(X; Y ) = H(Y ) − H(Y | X)

= H2 (p1 (1 − f )) − p1 H2 (f ).

(9.47)

We differentiate this expression with respect to p 1 , taking care not to confuse log 2 with log e : d 1 − p1 (1 − f ) I(X; Y ) = (1 − f ) log 2 − H2 (f ). dp1 p1 (1 − f )

(9.48)

Setting this derivative to zero and rearranging using skills developed in exercise 2.17 (p.36), we obtain: p∗1 (1 − f ) =

1 1+2

H2 (f )/(1−f )

,

(9.49)

so the optimal input distribution is p∗1

=

1

1/(1 − f )

1 + 2(H2 (f )/(1−f ))

Z BSC BEC

0.9

.

(9.50)

0.8 0.7 0.6

As the noise level f tends to 1, this expression tends to 1/e (as you can prove using L’Hˆopital’s rule). For all values of f, p∗1 is smaller than 1/2. A rough intuition for why input 1 is used less than input 0 is that when input 1 is used, the noisy channel injects entropy into the received string; whereas when input 0 is used, the noise has zero entropy. Solution to exercise 9.16 (p.155). The capacities of the three channels are shown in figure 9.11. For any f < 0.5, the BEC is the channel with highest capacity and the BSC the lowest. Solution to exercise 9.18 (p.155). ratio, given y, is a(y) = ln

The logarithm of the posterior probability

Q(y | x = 1, α, σ) αy P (x = 1 | y, α, σ) = ln =2 2. P (x = − 1 | y, α, σ) Q(y | x = − 1, α, σ) σ

(9.51)

0.5 0.4 0.3 0.2 0.1 0 0

0.1

0.2

0.3

0.4

0.5

0.6

0.7

0.8

0.9

1

Figure 9.11. Capacities of the Z channel, binary symmetric channel, and binary erasure channel.

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160

9 — Communication over a Noisy Channel

Using our skills picked up from exercise 2.17 (p.36), we rewrite this in the form 1 . (9.52) P (x = 1 | y, α, σ) = 1 + e−a(y) The optimal decoder selects the most probable hypothesis; this can be done simply by looking at the sign of a(y). If a(y) > 0 then decode as x ˆ = 1. The probability of error is pb =

Z

0

dy Q(y | x = 1, α, σ) =

−∞

Z

−xα

−∞

2

 xα  y 1 − dy √ . (9.53) e 2σ2 = Φ − σ 2πσ 2

Random coding Solution to exercise 9.20 (p.156). The probability that S = 24 people whose birthdays are drawn at random from A = 365 days all have distinct birthdays is A(A − 1)(A − 2) . . . (A − S + 1) . (9.54) AS The probability that two (or more) people share a birthday is one minus this quantity, which, for S = 24 and A = 365, is about 0.5. This exact way of answering the question is not very informative since it is not clear for what value of S the probability changes from being close to 0 to being close to 1. The number of pairs is S(S − 1)/2, and the probability that a particular pair shares a birthday is 1/A, so the expected number of collisions is S(S − 1) 1 . 2 A

(9.55)

This √ answer is more instructive. The expected number of collisions is tiny if √ S  A and big if S  A. We can also approximate the probability that all birthdays are distinct, for small S, thus: A(A − 1)(A − 2) . . . (A − S + 1) = (1)(1 − 1/A)(1 − 2/A) . . . (1 − (S −1)/A) AS ' exp(0) exp(−1/A) exp(−2/A) . . . exp(−(S −1)/A) (9.56) !   S−1 1 X S(S − 1)/2 ' exp − i = exp − . (9.57) A A i=1

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About Chapter 10 Before reading Chapter 10, you should have read Chapters 4 and 9. Exercise 9.14 (p.153) is especially recommended.

Cast of characters Q C XN C N x(s) s S = 2K K = log2 S s R = K/N sˆ

the noisy channel the capacity of the channel an ensemble used to create a random code a random code the length of the codewords a codeword, the sth in the code the number of a chosen codeword (mnemonic: the source selects s) the total number of codewords in the code the number of bits conveyed by the choice of one codeword from S, assuming it is chosen with uniform probability a binary representation of the number s the rate of the code, in bits per channel use (sometimes called R0 instead) the decoder’s guess of s

161

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10 The Noisy-Channel Coding Theorem 10.1 The theorem The theorem has three parts, two positive and one negative. The main positive result is the first.

pb 6

1. For every discrete memoryless channel, the channel capacity C = max I(X; Y ) PX

has the following property. For any  > 0 and R < C, for large enough N , there exists a code of length N and rate ≥ R and a decoding algorithm, such that the maximal probability of block error is < . 2. If a probability of bit error pb is acceptable, rates up to R(pb ) are achievable, where C R(pb ) = . (10.2) 1 − H2 (pb ) 3. For any pb , rates greater than R(pb ) are not achievable.

10.2 Jointly-typical sequences We formalize the intuitive preview of the last chapter. We will define codewords x(s) as coming from an ensemble X N , and consider the random selection of one codeword and a corresponding channel output y, thus defining a joint ensemble (XY ) N . We will use a typical-set decoder, which decodes a received signal y as s if x (s) and y are jointly typical, a term to be defined shortly. The proof will then centre on determining the probabilities (a) that the true input codeword is not jointly typical with the output sequence; and (b) that a false input codeword is jointly typical with the output. We will show that, for large N , both probabilities go to zero as long as there are fewer than 2N C codewords, and the ensemble X is the optimal input distribution. Joint typicality. A pair of sequences x, y of length N are defined to be jointly typical (to tolerance β) with respect to the distribution P (x, y) if 1 1 x is typical of P (x), i.e., log − H(X) < β, N P (x) 1 1 y is typical of P (y), i.e., log − H(Y ) < β, N P (y) 1 1 − H(X, Y ) < β. and x, y is typical of P (x, y), i.e., log N P (x, y) 162

R(pb )

(10.1) 1  C

2   

3 R

Figure 10.1. Portion of the R, pb plane to be proved achievable (1, 2) and not achievable (3).

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10.2: Jointly-typical sequences

163

The jointly-typical set JN β is the set of all jointly-typical sequence pairs of length N . Example. Here is a jointly-typical pair of length N = 100 for the ensemble P (x, y) in which P (x) has (p0 , p1 ) = (0.9, 0.1) and P (y | x) corresponds to a binary symmetric channel with noise level 0.2. x y

1111111111000000000000000000000000000000000000000000000000000000000000000000000000000000000000000000 0011111111000000000000000000000000000000000000000000000000000000000000000000000000111111111111111111

Notice that x has 10 1s, and so is typical of the probability P (x) (at any tolerance β); and y has 26 1s, so it is typical of P (y) (because P (y = 1) = 0.26); and x and y differ in 20 bits, which is the typical number of flips for this channel. Joint typicality theorem. Let x, y be drawn from the ensemble (XY ) N defined by N Y P (x, y) = P (xn , yn ). n=1

Then

1. the probability that x, y are jointly typical (to tolerance β) tends to 1 as N → ∞; 2. the number of jointly-typical sequences |J N β | is close to 2N H(X,Y ) . To be precise, |JN β | ≤ 2N (H(X,Y )+β) ; (10.3) 3. if x0 ∼ X N and y0 ∼ Y N , i.e., x0 and y0 are independent samples with the same marginal distribution as P (x, y), then the probability that (x0 , y0 ) lands in the jointly-typical set is about 2 −N I(X;Y ) . To be precise, P ((x0 , y0 ) ∈ JN β ) ≤ 2−N (I(X;Y )−3β) . (10.4)

Proof. The proof of parts 1 and 2 by the law of large numbers follows that of the source coding theorem in Chapter 4. For part 2, let the pair x, y play the role of x in the source coding theorem, replacing P (x) there by the probability distribution P (x, y). For the third part, P ((x0 , y0 ) ∈ JN β ) =

X

P (x)P (y)

(10.5)

(x,y)∈JN β

≤ |JN β | 2−N (H(X)−β) 2−N (H(Y )−β) ≤ 2

= 2

N (H(X,Y )+β)−N (H(X)+H(Y )−2β) −N (I(X;Y )−3β)

.

(10.6) (10.7)

2 (10.8)

A cartoon of the jointly-typical set is shown in figure 10.2. Two independent typical vectors are jointly typical with probability P ((x0 , y0 ) ∈ JN β ) ' 2−N (I(X;Y ))

(10.9)

because the total number of independent typical pairs is the area of the dashed rectangle, 2N H(X) 2N H(Y ) , and the number of jointly-typical pairs is roughly 2N H(X,Y ) , so the probability of hitting a jointly-typical pair is roughly 2N H(X,Y ) /2N H(X)+N H(Y ) = 2−N I(X;Y ) .

(10.10)

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164

10 — The Noisy-Channel Coding Theorem  6

 qq 6qq qq q qqq qq qq q q

AN X q q q q q q

2N H(X) q q q q q q

q q q q q q

AN Y

2N H(Y )

? ?

-

q q q q q q

q q q q q q

q q q q q q

2N H(X,Y ) dots q q q q q q

q q q q6 qq 6 q q q 2N H(Y |X) qqq ? qqq? qqq qq q qq qqqq qqq qqq qqq qqq qqq qqq qqq qqq  qqq N H(X|Y ) q q q 2 qqq qqq qqq qq qq

Figure 10.2. The jointly-typical set. The horizontal direction represents AN X , the set of all input strings of length N . The vertical direction represents AN Y , the set of all output strings of length N . The outer box contains all conceivable input–output pairs. Each dot represents a jointly-typical pair of sequences (x, y). The total number of jointly-typical sequences is about 2N H(X,Y ) .

10.3 Proof of the noisy-channel coding theorem Analogy Imagine that we wish to prove that there is a baby in a class of one hundred babies who weighs less than 10 kg. Individual babies are difficult to catch and weigh. Shannon’s method of solving the task is to scoop up all the babies and weigh them all at once on a big weighing machine. If we find that their average weight is smaller than 10 kg, there must exist at least one baby who weighs less than 10 kg – indeed there must be many! Shannon’s method isn’t guaranteed to reveal the existence of an underweight child, since it relies on there being a tiny number of elephants in the class. But if we use his method and get a total weight smaller than 1000 kg then our task is solved.

From skinny children to fantastic codes We wish to show that there exists a code and a decoder having small probability of error. Evaluating the probability of error of any particular coding and decoding system is not easy. Shannon’s innovation was this: instead of constructing a good coding and decoding system and evaluating its error probability, Shannon calculated the average probability of block error of all codes, and proved that this average is small. There must then exist individual codes that have small probability of block error.

Random coding and typical-set decoding Consider the following encoding–decoding system, whose rate is R 0 . 0

1. We fix P (x) and generate the S = 2N R codewords of a (N, N R0 ) =

Figure 10.3. Shannon’s method for proving one baby weighs less than 10 kg.

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10.3: Proof of the noisy-channel coding theorem x(3) x(1) x(2) x(4) qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq

165

x(3) x(1) x(2) x(4) qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq qq

ya yb

yd yc

(a)

- sˆ(ya ) = 0 - sˆ(yb ) = 3

- sˆ(yd ) = 0 - sˆ(yc ) = 4

(b)

(N, K) code C at random according to P (x) =

N Y

P (xn ).

(10.11)

n=1

A random code is shown schematically in figure 10.4a. 2. The code is known to both sender and receiver. 0

3. A message s is chosen from {1, 2, . . . , 2 N R }, and x(s) is transmitted. The received signal is y, with P (y | x(s) ) =

N Y

n=1

P (yn | x(s) n ).

(10.12)

4. The signal is decoded by typical-set decoding. Typical-set decoding. Decode y as sˆ if (x (ˆs) , y) are jointly typical and 0 there is no other s0 such that (x(s ) , y) are jointly typical; otherwise declare a failure (ˆ s = 0). This is not the optimal decoding algorithm, but it will be good enough, and easier to analyze. The typical-set decoder is illustrated in figure 10.4b. 5. A decoding error occurs if sˆ 6= s. There are three probabilities of error that we can distinguish. First, there is the probability of block error for a particular code C, that is, pB (C) ≡ P (ˆ s 6= s | C).

(10.13)

This is a difficult quantity to evaluate for any given code. Second, there is the average over all codes of this block error probability, X hpB i ≡ P (ˆ s 6= s | C)P (C). (10.14) C

Figure 10.4. (a) A random code. (b) Example decodings by the typical set decoder. A sequence that is not jointly typical with any of the codewords, such as ya , is decoded as sˆ = 0. A sequence that is jointly typical with codeword x(3) alone, yb , is decoded as sˆ = 3. Similarly, yc is decoded as sˆ = 4. A sequence that is jointly typical with more than one codeword, such as yd , is decoded as sˆ = 0.

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166

10 — The Noisy-Channel Coding Theorem

Fortunately, this quantity is much easier to evaluate than the first quantity P (ˆ s 6= s | C). Third, the maximal block error probability of a code C, pBM (C) ≡ max P (ˆ s 6= s | s, C),

(10.15)

s

is the quantity we are most interested in: we wish to show that there exists a code C with the required rate whose maximal block error probability is small. We will get to this result by first finding the average block error probability, hpB i. Once we have shown that this can be made smaller than a desired small number, we immediately deduce that there must exist at least one code C whose block error probability is also less than this small number. Finally, we show that this code, whose block error probability is satisfactorily small but whose maximal block error probability is unknown (and could conceivably be enormous), can be modified to make a code of slightly smaller rate whose maximal block error probability is also guaranteed to be small. We modify the code by throwing away the worst 50% of its codewords. We therefore now embark on finding the average probability of block error.

Probability of error of typical-set decoder There are two sources of error when we use typical-set decoding. Either (a) the output y is not jointly typical with the transmitted codeword x (s) , or (b) there is some other codeword in C that is jointly typical with y. By the symmetry of the code construction, the average probability of error averaged over all codes does not depend on the selected value of s; we can assume without loss of generality that s = 1. (a) The probability that the input x (1) and the output y are not jointly typical vanishes, by the joint typicality theorem’s first part (p.163). We give a name, δ, to the upper bound on this probability, satisfying δ → 0 as N → ∞; for any desired δ, we can find a blocklength N (δ) such that the P ((x (1) , y) 6∈ JN β ) ≤ δ. 0 (b) The probability that x(s ) and y are jointly typical, for a given s 0 6= 1 0 −N (I(X;Y )−3β) is ≤ 2 , by part 3. And there are (2N R − 1) rival values of s0 to worry about. Thus the average probability of error hp B i satisfies: 0

hpB i ≤ δ +

NR 2X

2−N (I(X;Y )−3β)

s0 =2 −N (I(X;Y )−R0 −3β)

≤ δ+2

.

(10.16) (10.17)

The inequality (10.16) that bounds a total probability of error PTOT by the sum of the probabilities Ps0 of all sorts of events s0 each of which is sufficient to cause error, PTOT ≤ P1 + P2 + · · · ,

is called a union bound. It is only an equality if the different events that cause error never occur at the same time as each other.

The average probability of error (10.17) can be made < 2δ by increasing N if R0 < I(X; Y ) − 3β.

(10.18)

We are almost there. We make three modifications: 1. We choose P (x) in the proof to be the optimal input distribution of the channel. Then the condition R 0 < I(X; Y ) − 3β becomes R0 < C − 3β.

hpB i is just the probability that there is a decoding error at step 5 of the five-step process on the previous page.

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10.4: Communication (with errors) above capacity

(a) A random code . . .



167

(b) expurgated

Figure 10.5. How expurgation works. (a) In a typical random code, a small fraction of the codewords are involved in collisions – pairs of codewords are sufficiently close to each other that the probability of error when either codeword is transmitted is not tiny. We obtain a new code from a random code by deleting all these confusable codewords. (b) The resulting code has slightly fewer codewords, so has a slightly lower rate, and its maximal probability of error is greatly reduced.

2. Since the average probability of error over all codes is < 2δ, there must exist a code with mean probability of block error p B (C) < 2δ. 3. To show that not only the average but also the maximal probability of error, pBM , can be made small, we modify this code by throwing away the worst half of the codewords – the ones most likely to produce errors. Those that remain must all have conditional probability of error less than 4δ. We use these remaining codewords to define a new code. This 0 new code has 2N R −1 codewords, i.e., we have reduced the rate from R 0 to R0 − 1/N (a negligible reduction, if N is large), and achieved p BM < 4δ. This trick is called expurgation (figure 10.5). The resulting code may not be the best code of its rate and length, but it is still good enough to prove the noisy-channel coding theorem, which is what we are trying to do here. In conclusion, we can ‘construct’ a code of rate R 0 − 1/N , where R0 < C − 3β, with maximal probability of error < 4δ. We obtain the theorem as stated by setting R0 = (R + C)/2, δ = /4, β < (C − R 0 )/3, and N sufficiently large for the remaining conditions to hold. The theorem’s first part is thus proved. 2

10.4 Communication (with errors) above capacity We have proved, for any discrete memoryless channel, the achievability of a portion of the R, pb plane shown in figure 10.6. We have shown that we can turn any noisy channel into an essentially noiseless binary channel with rate up to C bits per cycle. We now extend the right-hand boundary of the region of achievability at non-zero error probabilities. [This is called rate-distortion theory.] We do this with a new trick. Since we know we can make the noisy channel into a perfect channel with a smaller rate, it is sufficient to consider communication with errors over a noiseless channel. How fast can we communicate over a noiseless channel, if we are allowed to make errors? Consider a noiseless binary channel, and assume that we force communication at a rate greater than its capacity of 1 bit. For example, if we require the sender to attempt to communicate at R = 2 bits per cycle then he must effectively throw away half of the information. What is the best way to do this if the aim is to achieve the smallest possible probability of bit error? One simple strategy is to communicate a fraction 1/R of the source bits, and ignore the rest. The receiver guesses the missing fraction 1 − 1/R at random, and

pb 6

achievable -

C

R

Figure 10.6. Portion of the R, pb plane proved achievable in the first part of the theorem. [We’ve proved that the maximal probability of block error pBM can be made arbitrarily small, so the same goes for the bit error probability pb , which must be smaller than pBM .]

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168

10 — The Noisy-Channel Coding Theorem 0.3

the average probability of bit error is 0.25

1 pb = (1 − 1/R). 2

(10.19)

pb

Optimum Simple

0.2

0.15

The curve corresponding to this strategy is shown by the dashed line in figure 10.7. We can do better than this (in terms of minimizing p b ) by spreading out the risk of corruption evenly among all the bits. In fact, we can achieve pb = H2−1 (1 − 1/R), which is shown by the solid curve in figure 10.7. So, how can this optimum be achieved? We reuse a tool that we just developed, namely the (N, K) code for a noisy channel, and we turn it on its head, using the decoder to define a lossy compressor. Specifically, we take an excellent (N, K) code for the binary symmetric channel. Assume that such a code has a rate R 0 = K/N , and that it is capable of correcting errors introduced by a binary symmetric channel whose transition probability is q. Asymptotically, rate-R 0 codes exist that have R0 ' 1 − H2 (q). Recall that, if we attach one of these capacity-achieving codes of length N to a binary symmetric channel then (a) the probability distribution over the outputs is close to uniform, since the entropy of the output is equal to the entropy of the source (N R 0 ) plus the entropy of the noise (N H2 (q)), and (b) the optimal decoder of the code, in this situation, typically maps a received vector of length N to a transmitted vector differing in qN bits from the received vector. We take the signal that we wish to send, and chop it into blocks of length N (yes, N , not K). We pass each block through the decoder, and obtain a shorter signal of length K bits, which we communicate over the noiseless channel. To decode the transmission, we pass the K bit message to the encoder of the original code. The reconstituted message will now differ from the original message in some of its bits – typically qN of them. So the probability of bit error will be pb = q. The rate of this lossy compressor is R = N/K = 1/R 0 = 1/(1 − H2 (pb )). Now, attaching this lossy compressor to our capacity-C error-free communicator, we have proved the achievability of communication up to the curve (pb , R) defined by: C . 2 (10.20) R= 1 − H2 (pb )

0.1 0.05 0 0

0.5

1

1.5

Figure 10.7. A simple bound on achievable points (R, pb ), and Shannon’s bound.

For further reading about rate-distortion theory, see Gallager (1968), p. 451, or McEliece (2002), p. 75.

10.5 The non-achievable region (part 3 of the theorem) The source, encoder, noisy channel and decoder define a Markov chain: P (s, x, y, sˆ) = P (s)P (x | s)P (y | x)P (ˆ s | y).

(10.21)

The data processing inequality (exercise 8.9, p.141) must apply to this chain: I(s; sˆ) ≤ I(x; y). Furthermore, by the definition of channel capacity, I(x; y) ≤ N C, so I(s; sˆ) ≤ N C. Assume that a system achieves a rate R and a bit error probability p b ; then the mutual information I(s; sˆ) is ≥ N R(1 − H 2 (pb )). But I(s; sˆ) > N C is not achievable, so R > 1−HC (p ) is not achievable. 2 2

b

Exercise 10.1.[3 ] Fill in the details in the preceding argument. If the bit errors between sˆ and s are independent then we have I(s; sˆ) = N R(1−H 2 (pb )).

2

R

s → x → y → sˆ

2.5

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10.6: Computing capacity

169

What if we have complex correlations among those bit errors? Why does the inequality I(s; sˆ) ≥ N R(1 − H2 (pb )) hold?

10.6 Computing capacity We have proved that the capacity of a channel is the maximum rate at which reliable communication can be achieved. How can we compute the capacity of a given discrete memoryless channel? We need to find its optimal input distribution. In general we can find the optimal input distribution by a computer search, making use of the derivative of the mutual information with respect to the input probabilities.

Sections 10.6–10.8 contain advanced material. The first-time reader is encouraged to skip to section 10.9 (p.172).

. Exercise 10.2.[2 ] Find the derivative of I(X; Y ) with respect to the input probability pi , ∂I(X; Y )/∂pi , for a channel with conditional probabilities Q j|i . Exercise 10.3.[2 ] Show that I(X; Y ) is a concave _ function of the input probability vector p. Since I(X; Y ) is concave _ in the input distribution p, any probability distribution p at which I(X; Y ) is stationary must be a global maximum of I(X; Y ). So it is tempting to put the derivative of I(X; Y ) into a routine that finds a local maximum of I(X; Y ), that is, an input distribution P (x) such that ∂I(X; Y ) = λ for all i, ∂pi

(10.22)

P where λ is a Lagrange multiplier associated with the constraint i pi = 1. However, this approach may fail to find the right answer, because I(X; Y ) might be maximized by a distribution that has p i = 0 for some inputs. A simple example is given by the ternary confusion channel. Ternary confusion channel. AX = {0, ?, 1}. AY = {0, 1}. 0 -0  P (y = 0 | x = 0) = 1 ; P (y = 0 | x = ?) = 1/2 ; P (y = 0 | x = 1) = 0 ; ? @ R @ 1 1 P (y = 1 | x = 0) = 0 ; P (y = 1 | x = ?) = 1/2 ; P (y = 1 | x = 1) = 1. Whenever the input ? is used, the output is random; the other inputs are reliable inputs. The maximum information rate of 1 bit is achieved by making no use of the input ?. . Exercise 10.4.[2, p.173] Sketch the mutual information for this channel as a function of the input distribution p. Pick a convenient two-dimensional representation of p. The optimization routine must therefore take account of the possibility that, as we go up hill on I(X; Y ), we may run into the inequality constraints p i ≥ 0. . Exercise 10.5.[2, p.174] Describe the condition, similar to equation (10.22), that is satisfied at a point where I(X; Y ) is maximized, and describe a computer program for finding the capacity of a channel.

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170

10 — The Noisy-Channel Coding Theorem

Results that may help in finding the optimal input distribution 1. All outputs must be used. 2. I(X; Y ) is a convex ^ function of the channel parameters. 3. There may be several optimal input distributions, but they all look the same at the output. . Exercise 10.6.[2 ] Prove that no output y is unused by an optimal input distribution, unless it is unreachable, that is, has Q(y | x) = 0 for all x. Exercise 10.7.[2 ] Prove that I(X; Y ) is a convex ^ function of Q(y | x). Exercise 10.8.[2 ] Prove that all optimal input distributions P of a channel have the same output probability distribution P (y) = x P (x)Q(y | x).

These results, along with the fact that I(X; Y ) is a concave _ function of the input probability vector p, prove the validity of the symmetry argument that we have used when finding the capacity of symmetric channels. If a channel is invariant under a group of symmetry operations – for example, interchanging the input symbols and interchanging the output symbols – then, given any optimal input distribution that is not symmetric, i.e., is not invariant under these operations, we can create another input distribution by averaging together this optimal input distribution and all its permuted forms that we can make by applying the symmetry operations to the original optimal input distribution. The permuted distributions must have the same I(X; Y ) as the original, by symmetry, so the new input distribution created by averaging must have I(X; Y ) bigger than or equal to that of the original distribution, because of the concavity of I.

Symmetric channels In order to use symmetry arguments, it will help to have a definition of a symmetric channel. I like Gallager’s (1968) definition. A discrete memoryless channel is a symmetric channel if the set of outputs can be partitioned into subsets in such a way that for each subset the matrix of transition probabilities has the property that each row (if more than 1) is a permutation of each other row and each column is a permutation of each other column. Example 10.9. This channel P (y = 0 | x = 0) = 0.7 ; P (y = 0 | x = 1) = 0.1 ; P (y = ? | x = 0) = 0.2 ; P (y = ? | x = 1) = 0.2 ; P (y = 1 | x = 0) = 0.1 ; P (y = 1 | x = 1) = 0.7.

(10.23)

is a symmetric channel because its outputs can be partitioned into (0, 1) and ?, so that the matrix can be rewritten: P (y = 0 | x = 0) = 0.7 ; P (y = 1 | x = 0) = 0.1 ;

P (y = 0 | x = 1) = 0.1 ; P (y = 1 | x = 1) = 0.7 ;

P (y = ? | x = 0) = 0.2 ;

P (y = ? | x = 1) = 0.2.

(10.24)

Reminder: The term ‘convex ^’ means ‘convex’, and the term ‘concave _’ means ‘concave’; the little smile and frown symbols are included simply to remind you what convex and concave mean.

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10.7: Other coding theorems

171

Symmetry is a useful property because, as we will see in a later chapter, communication at capacity can be achieved over symmetric channels by linear codes. Exercise 10.10. [2 ] Prove that for a symmetric channel with any number of inputs, the uniform distribution over the inputs is an optimal input distribution. . Exercise 10.11. [2, p.174] Are there channels that are not symmetric whose optimal input distributions are uniform? Find one, or prove there are none.

10.7 Other coding theorems The noisy-channel coding theorem that we proved in this chapter is quite general, applying to any discrete memoryless channel; but it is not very specific. The theorem only says that reliable communication with error probability  and rate R can be achieved by using codes with sufficiently large blocklength N . The theorem does not say how large N needs to be to achieve given values of R and . Presumably, the smaller  is and the closer R is to C, the larger N has to be.

Noisy-channel coding theorem – version with explicit N -dependence For a discrete memoryless channel, a blocklength N and a rate R, there exist block codes of length N whose average probability of error satisfies: pB ≤ exp [−N Er (R)] (10.25) where Er (R) is the random-coding exponent of the channel, a convex ^, decreasing, positive function of R for 0 ≤ R < C. The random-coding exponent is also known as the reliability function. [By an expurgation argument it can also be shown that there exist block codes for which the maximal probability of error p BM is also exponentially small in N .] The definition of Er (R) is given in Gallager (1968), p. 139. E r (R) approaches zero as R → C; the typical behaviour of this function is illustrated in figure 10.8. The computation of the random-coding exponent for interesting channels is a challenging task on which much effort has been expended. Even for simple channels like the binary symmetric channel, there is no simple expression for Er (R).

Lower bounds on the error probability as a function of blocklength The theorem stated above asserts that there are codes with p B smaller than exp [−N Er (R)]. But how small can the error probability be? Could it be much smaller? For any code with blocklength N on a discrete memoryless channel, the probability of error assuming all source messages are used with equal probability satisfies pB

exp[−N Esp (R)],

(10.26)

Er (R)

C R Figure 10.8. A typical random-coding exponent.

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172

10 — The Noisy-Channel Coding Theorem where the function Esp (R), the sphere-packing exponent of the channel, is a convex ^, decreasing, positive function of R for 0 ≤ R < C.

For a precise statement of this result and further references, see Gallager (1968), p. 157.

10.8 Noisy-channel coding theorems and coding practice Imagine a customer who wants to buy an error-correcting code and decoder for a noisy channel. The results described above allow us to offer the following service: if he tells us the properties of his channel, the desired rate R and the desired error probability pB , we can, after working out the relevant functions C, Er (R), and Esp (R), advise him that there exists a solution to his problem using a particular blocklength N ; indeed that almost any randomly chosen code with that blocklength should do the job. Unfortunately we have not found out how to implement these encoders and decoders in practice; the cost of implementing the encoder and decoder for a random code with large N would be exponentially large in N . Furthermore, for practical purposes, the customer is unlikely to know exactly what channel he is dealing with. So Berlekamp (1980) suggests that the sensible way to approach error-correction is to design encoding-decoding systems and plot their performance on a variety of idealized channels as a function of the channel’s noise level. These charts (one of which is illustrated on page 568) can then be shown to the customer, who can choose among the systems on offer without having to specify what he really thinks his channel is like. With this attitude to the practical problem, the importance of the functions Er (R) and Esp (R) is diminished.

10.9 Further exercises Exercise 10.12. [2 ] A binary erasure channel with input x and output y has transition probability matrix:   1−q 0 0 -0 @   R @? q q Q=  0 1−q 1 -1 Find the mutual information I(X; Y ) between the input and output for general input distribution {p0 , p1 }, and show that the capacity of this channel is C = 1 − q bits. A Z channel has transition probability matrix: Q=



1 q 0 1−q



0-0    1 1

Show that, using a (2, 1) code, two uses of a Z channel can be made to emulate one use of an erasure channel, and state the erasure probability of that erasure channel. Hence show that the capacity of the Z channel, CZ , satisfies CZ ≥ 12 (1 − q) bits. Explain why the result CZ ≥ equality.

1 2 (1

− q) is an inequality rather than an

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10.10: Solutions

173

Exercise 10.13. [3, p.174] A transatlantic cable contains N = 20 indistinguishable electrical wires. You have the job of figuring out which wire is which, that is, to create a consistent labelling of the wires at each end. Your only tools are the ability to connect wires to each other in groups of two or more, and to test for connectedness with a continuity tester. What is the smallest number of transatlantic trips you need to make, and how do you do it? How would you solve the problem for larger N such as N = 1000? As an illustration, if N were 3 then the task can be solved in two steps by labelling one wire at one end a, connecting the other two together, crossing the Atlantic, measuring which two wires are connected, labelling them b and c and the unconnected one a, then connecting b to a and returning across the Atlantic, whereupon on disconnecting b from c, the identities of b and c can be deduced. This problem can be solved by persistent search, but the reason it is posed in this chapter is that it can also be solved by a greedy approach based on maximizing the acquired information. Let the unknown permutation of wires be x. Having chosen a set of connections of wires C at one end, you can then make measurements at the other end, and these measurements y convey information about x. How much? And for what set of connections is the information that y conveys about x maximized?

10.10 Solutions Solution to exercise 10.4 (p.169). the mutual information is

If the input distribution is p = (p 0 , p? , p1 ),

I(X; Y ) = H(Y ) − H(Y |X) = H2 (p0 + p? /2) − p? .

- 0

0

(10.27)

We can build a good sketch of this function in two ways: by careful inspection of the function, or by looking at special cases. For the plots, the two-dimensional representation of p I will use has p 0 and p1 as the independent variables, so that p = (p 0 , p? , p1 ) = (p0 , (1−p0 −p1 ), p1 ).

 1/2

?

@ 1/2 @ R @ 1 - 1

By inspection. If we use the quantities p ∗ ≡ p0 + p? /2 and p? as our two degrees of freedom, the mutual information becomes very simple: I(X; Y ) = H2 (p∗ ) − p? . Converting back to p0 = p∗ − p? /2 and p1 = 1 − p∗ − p? /2, we obtain the sketch shown at the left below. This function is like a tunnel rising up the direction of increasing p 0 and p1 . To obtain the required plot of I(X; Y ) we have to strip away the parts of this tunnel that live outside the feasible simplex of probabilities; we do this by redrawing the surface, showing only the parts where p0 > 0 and p1 > 0. A full plot of the function is shown at the right. 1

1

0.5 0

0.5 1 p1 0.5

-0.5 -1 -0.5

0 -0.5 0

0.5

p0

1

1 0.5 0 0

0 0.5

p0

1

p1

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174

10 — The Noisy-Channel Coding Theorem

Special cases. In the special case p ? = 0, the channel is a noiseless binary channel, and I(X; Y ) = H2 (p0 ). In the special case p0 = p1 , the term H2 (p0 + p? /2) is equal to 1, so I(X; Y ) = 1 − p? . In the special case p0 = 0, the channel is a Z channel with error probability 0.5. We know how to sketch that, from the previous chapter (figure 9.3). These special cases allow us to construct the skeleton shown in figure 10.9. Solution to exercise 10.5 (p.169). Necessary and sufficient conditions for p to maximize I(X; Y ) are ) ∂I(X;Y ) = λ and pi > 0 ∂pi for all i, (10.28) ∂I(X;Y ) ≤ λ and pi = 0 ∂pi where λ is a constant related to the capacity by C = λ + log 2 e. This result can be used in a computer program that evaluates the derivatives, and increments and decrements the probabilities p i in proportion to the differences between those derivatives. This result is also useful for lazy human capacity-finders who are good guessers. Having guessed the optimal input distribution, one can simply confirm that equation (10.28) holds. Solution to exercise 10.11 (p.171). We certainly expect nonsymmetric channels with uniform optimal input distributions to exist, since when inventing a channel we have I(J − 1) degrees of freedom whereas the optimal input distribution is just (I − 1)-dimensional; so in the I(J −1)-dimensional space of perturbations around a symmetric channel, we expect there to be a subspace of perturbations of dimension I(J − 1) − (I − 1) = I(J − 2) + 1 that leave the optimal input distribution unchanged. Here is an explicit example, a bit like a Z channel.   0.9585 0.0415 0.35 0.0  0.0415 0.9585 0.0 0.35   Q= (10.29)  0 0 0.65 0  0 0 0 0.65

Solution to exercise 10.13 (p.173). The labelling problem can be solved for any N > 2 with just two trips, one each way across the Atlantic. The key step in the information-theoretic approach to this problem is to write down the information content of one partition, the combinatorial object that is the connecting together of subsets of wires. If N wires are grouped together into g1 subsets of size 1, g2 subsets of size 2, . . . , then the number of such partitions is N! Ω= Y , (10.30) (r!)gr gr ! r

and the information content of one such partition is the log of this quantity. In a greedy strategy we choose the first partition to maximize this information content. One game we can play is to maximize this information content with respect to the quantities gr , treated as real numbers, subject to the constraint P g r r r = N . Introducing a Lagrange multiplier λ for the constraint, the derivative is ! X ∂ log Ω + λ gr r = − log r! − log gr + λr, (10.31) ∂gr r

1

p1

0.5

1 0.5 0 0

0.5

p0

0 1

Figure 10.9. Skeleton of the mutual information for the ternary confusion channel.

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10.10: Solutions

175

which, when set to zero, leads to the rather nice expression gr =

eλr ; r!

(10.32)

the optimal gr is proportional to a Poisson distribution! We P can solve for the Lagrange multiplier by plugging gr into the constraint r gr r = N , which gives the implicit equation N = µ eµ , (10.33)

(a)

1

10

100

1000

2.5 2



where µ ≡ is a convenient reparameterization of the Lagrange multiplier. Figure 10.10a shows a graph of µ(N ); figure 10.10b shows the deduced noninteger assignments gr when µ = 2.2, and nearby integers gr = {1, 2, 2, 1, 1} that motivate setting the first partition to (a)(bc)(de)(fgh)(ijk)(lmno)(pqrst). This partition produces a random partition at the other end, which has an information content of log Ω = 40.4 bits, which is a lot more than half the total information content we need to acquire to infer the transatlantic permutation, log 20! ' 61 bits. [In contrast, if all the wires are joined together in pairs, the information content generated is only about 29 bits.] How to choose the second partition is left to the reader. A Shannonesque approach is appropriate, picking a random partition at the other end, using the same {g r }; you need to ensure the two partitions are as unlike each other as possible. If N 6= 2, 5 or 9, then the labelling problem has solutions that are particularly simple to implement, called Knowlton–Graham partitions: partition {1, . . . , N } into disjoint sets in two ways A and B, subject to the condition that at most one element appears both in an A set of cardinality j and in a B set of cardinality k, for each j and k (Graham, 1966; Graham and Knowlton, 1968).

5.5 5 4.5 4 3.5 3 2.5 2 1.5 1 0.5

1.5 1 0.5

(b)

0 1

2

3

4

5

6

7

8

9

10

Figure 10.10. Approximate solution of the cable-labelling problem using Lagrange multipliers. (a) The parameter µ as a function of N ; the value µ(20) = 2.2 is highlighted. (b) Non-integer values of the function r gr = µ /r! are shown by lines and integer values of gr motivated by those non-integer values are shown by crosses.

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About Chapter 11 Before reading Chapter 11, you should have read Chapters 9 and 10. You will also need to be familiar with the Gaussian distribution. One-dimensional Gaussian distribution. If a random variable y is Gaussian and has mean µ and variance σ 2 , which we write: y ∼ Normal(µ, σ 2 ), or P (y) = Normal(y; µ, σ 2 ),

(11.1)

then the distribution of y is:   1 P (y | µ, σ 2 ) = √ exp −(y − µ)2 /2σ 2 . 2 2πσ

(11.2)

[I use the symbol P for both probability densities and probabilities.] The inverse-variance τ ≡ 1/σ 2 is sometimes called the precision of the Gaussian distribution. Multi-dimensional Gaussian distribution. If y = (y 1 , y2 , . . . , yN ) has a multivariate Gaussian distribution, then   1 1 exp − (y − x)TA(y − x) , (11.3) P (y | x, A) = Z(A) 2 where x is the mean of the distribution, A is the inverse of the variance–covariance matrix, and the normalizing constant is Z(A) = (det(A/2π))−1/2 . This distribution has the property that the variance Σ ii of yi , and the covariance Σij of yi and yj are given by Σij ≡ E [(yi − y¯i )(yj − y¯j )] = A−1 ij ,

(11.4)

where A−1 is the inverse of the matrix A. The marginal distribution P (yi ) of one component yi is Gaussian; the joint marginal distribution of any subset of the components is multivariate-Gaussian; and the conditional density of any subset, given the values of another subset, for example, P (y i | yj ), is also Gaussian.

176

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11 Error-Correcting Codes & Real Channels The noisy-channel coding theorem that we have proved shows that there exist reliable error-correcting codes for any noisy channel. In this chapter we address two questions. First, many practical channels have real, rather than discrete, inputs and outputs. What can Shannon tell us about these continuous channels? And how should digital signals be mapped into analogue waveforms, and vice versa? Second, how are practical error-correcting codes made, and what is achieved in practice, relative to the possibilities proved by Shannon?

11.1 The Gaussian channel The most popular model of a real-input, real-output channel is the Gaussian channel. The Gaussian channel has a real input x and a real output y. The conditional distribution of y given x is a Gaussian distribution:   1 P (y | x) = √ exp −(y − x)2 /2σ 2 . 2πσ 2

(11.5)

This channel has a continuous input and output but is discrete in time. We will show below that certain continuous-time channels are equivalent to the discrete-time Gaussian channel. This channel is sometimes called the additive white Gaussian noise (AWGN) channel. As with discrete channels, we will discuss what rate of error-free information communication can be achieved over this channel.

Motivation in terms of a continuous-time channel Consider a physical (electrical, say) channel with inputs and outputs that are continuous in time. We put in x(t), and out comes y(t) = x(t) + n(t). Our transmission has a power cost. The average power of a transmission of length T may be constrained thus: Z

T 0

dt [x(t)]2 /T ≤ P.

(11.6)

The received signal is assumed to differ from x(t) by additive noise n(t) (for example Johnson noise), which we will model as white Gaussian noise. The magnitude of this noise is quantified by the noise spectral density, N 0 . 177

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178

11 — Error-Correcting Codes and Real Channels

How could such a channel be used to communicate information? Consider transmitting a set of N real numbers {x n }N n=1 in a signal of duration T made up of a weighted combination of orthonormal basis functions φ n (t), N X

φ1 (t)

(11.7)

φ2 (t)

dt φn (t)φm (t) = δnm . The receiver can then compute the scalars: Z T Z T dt φn (t)n(t) (11.8) dt φn (t)y(t) = xn + yn ≡

φ3 (t)

x(t) =

xn φn (t),

n=1

where

RT 0

0

0

≡ xn + nn

(11.9)

for n = 1 . . . N . If there were no noise, then y n would equal xn . The white Gaussian noise n(t) adds scalar noise n n to the estimate yn . This noise is Gaussian: nn ∼ Normal(0, N0 /2), (11.10)

where N0 is the spectral density introduced above. Thus a continuous channel used in this way is equivalent Rto the Gaussian channel defined at equaT tion (11.5). The power constraint 0 dt [x(t)]2 ≤ P T defines a constraint on the signal amplitudes xn , X PT x2n ≤ P T ⇒ x2n ≤ . (11.11) N n

x(t)

Figure 11.1. Three basis functions, and a weighted combination of PN them, x(t) = n=1 xn φn (t), with x1 = 0.4, x2 = − 0.2, and x3 = 0.1.

Before returning to the Gaussian channel, we define the bandwidth (measured in Hertz) of the continuous channel to be:

N max , (11.12) 2T max where N is the maximum number of orthonormal functions that can be produced in an interval of length T . This definition can be motivated by imagining creating a band-limited signal of duration T from orthonormal cosine and sine curves of maximum frequency W . The number of orthonormal functions is N max = 2W T . This definition relates to the Nyquist sampling theorem: if the highest frequency present in a signal is W , then the signal can be fully determined from its values at a series of discrete sample points separated by the Nyquist interval ∆t = 1/2W seconds. So the use of a real continuous channel with bandwidth W , noise spectral density N0 , and power P is equivalent to N/T = 2W uses per second of a Gaussian channel with noise level σ 2 = N0 /2 and subject to the signal power constraint x2n ≤ P/2W . W =

Definition of Eb /N0 Imagine that the Gaussian channel y n = xn + nn is used with an encoding system to transmit binary source bits at a rate of R bits per channel use. How can we compare two encoding systems that have different rates of communication R and that use different powers x 2n ? Transmitting at a large rate R is good; using small power is good too. It is conventional to measure the rate-compensated signal-to-noise ratio by the ratio of the power per source bit E b = x2n /R to the noise spectral density N0 : x2 Eb /N0 = 2n . (11.13) 2σ R Eb /N0 is one of the measures used to compare coding schemes for Gaussian channels.

Eb /N0 is dimensionless, but it is usually reported in the units of decibels; the value given is 10 log10 Eb /N0 .

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11.2: Inferring the input to a real channel

179

11.2 Inferring the input to a real channel ‘The best detection of pulses’ In 1944 Shannon wrote a memorandum (Shannon, 1993) on the problem of best differentiating between two types of pulses of known shape, represented by vectors x0 and x1 , given that one of them has been transmitted over a noisy channel. This is a pattern recognition problem. It is assumed that the noise is Gaussian with probability density 

P (n) = det



A 2π

1/2





1 exp − nTAn , 2

(11.14)

where A is the inverse of the variance–covariance matrix of the noise, a symmetric and positive-definite matrix. (If A is a multiple of the identity matrix, I/σ 2 , then the noise is ‘white’. For more general A, the noise is ‘coloured’.) The probability of the received vector y given that the source signal was s (either zero or one) is then 

P (y | s) = det



A 2π

1/2





1 exp − (y − xs )TA(y − xs ) . 2

(11.15)

x0 x1 y

Figure 11.2. Two pulses x0 and x1 , represented as 31-dimensional vectors, and a noisy version of one of them, y.

The optimal detector is based on the posterior probability ratio: P (s = 1 | y) P (y | s = 1) P (s = 1) = (11.16) P (s = 0 | y) P (y | s = 0) P (s = 0)   1 1 P (s = 1) T T = exp − (y − x1 ) A(y − x1 ) + (y − x0 ) A(y − x0 ) + ln 2 2 P (s = 0) T = exp (y A(x1 − x0 ) + θ) , (11.17) where θ is a constant independent of the received vector y, 1 P (s = 1) 1 θ = − xT1 Ax1 + xT0 Ax0 + ln . 2 2 P (s = 0)

(11.18)

If the detector is forced to make a decision (i.e., guess either s = 1 or s = 0) then the decision that minimizes the probability of error is to guess the most probable hypothesis. We can write the optimal decision in terms of a discriminant function: a(y) ≡ yTA(x1 − x0 ) + θ (11.19) with the decisions

a(y) > 0 → guess s = 1 a(y) < 0 → guess s = 0 a(y) = 0 → guess either.

(11.20)

Notice that a(y) is a linear function of the received vector, a(y) = wTy + θ,

(11.21)

where w ≡ A(x1 − x0 ).

11.3 Capacity of Gaussian channel Until now we have measured the joint, marginal, and conditional entropy of discrete variables only. In order to define the information conveyed by continuous variables, there are two issues we must address – the infinite length of the real line, and the infinite precision of real numbers.

w

Figure 11.3. The weight vector w ∝ x1 − x0 that is used to discriminate between x0 and x1 .

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11 — Error-Correcting Codes and Real Channels

Infinite inputs How much information can we convey in one use of a Gaussian channel? If we are allowed to put any real number x into the Gaussian channel, we could communicate an enormous string of N digits d 1 d2 d3 . . . dN by setting x = d1 d2 d3 . . . dN 000 . . . 000. The amount of error-free information conveyed in just a single transmission could be made arbitrarily large by increasing N , and the communication could be made arbitrarily reliable by increasing the number of zeroes at the end of x. There is usually some power cost associated with large inputs, however, not to mention practical limits in the dynamic range acceptable to a receiver. It is therefore conventional to introduce a cost function v(x) for every input x, and constrain codes to have an average cost v¯ less than or equal to some maximum value. A generalized channel coding theorem, including a cost function for the inputs, can be proved – see McEliece (1977). The result is a channel capacity C(¯ v ) that is a function of the permitted cost. For the Gaussian channel we will assume a cost v(x) = x2

(a)

g  (b)

(11.22)

such that the ‘average power’ x2 of the input is constrained. We motivated this cost function above in the case of real electrical channels in which the physical power consumption is indeed quadratic in x. The constraint x 2 = v¯ makes it impossible to communicate infinite information in one use of the Gaussian channel.

Infinite precision It is tempting to define joint, marginal, and conditional entropies for real variables simply by replacing summations by integrals, but this is not a well defined operation. As we discretize an interval into smaller and smaller divisions, the entropy of the discrete distribution diverges (as the logarithm of the granularity) (figure 11.4). Also, it is not permissible to take the logarithm of a dimensional quantity such as a probability density P (x) (whose dimensions are [x]−1 ). There is one information measure, however, that has a well-behaved limit, namely the mutual information – and this is the one that really matters, since it measures how much information one variable conveys about another. In the discrete case, X P (x, y) I(X; Y ) = P (x, y) log . (11.23) P (x)P (y) x,y Now because the argument of the log is a ratio of two probabilities over the same space, it is OK to have P (x, y), P (x) and P (y) be probability densities and replace the sum by an integral: Z P (x, y) (11.24) I(X; Y ) = dx dy P (x, y) log P (x)P (y) Z P (y | x) = dx dy P (x)P (y | x) log . (11.25) P (y) We can now ask these questions for the Gaussian channel: (a) what probability distribution P (x) maximizes the mutual information (subject to the constraint x2 = v)? and (b) does the maximal mutual information still measure the maximum error-free communication rate of this real channel, as it did for the discrete channel?

.. .

Figure 11.4. (a) A probability density P (x). Question: can we define the ‘entropy’ of this density? (b) We could evaluate the entropies of a sequence of probability distributions with decreasing grain-size g, but these entropies tend to Z 1 dx, which is not P (x) log P (x)g independent of g: the entropy goes up by one bit for every halving of g. Z 1 dx is an illegal P (x) log P (x) integral.

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11.3: Capacity of Gaussian channel

181

Exercise 11.1.[3, p.189] Prove that the probability distribution P (x) that maximizes the mutual information (subject to the constraint x 2 = v) is a Gaussian distribution of mean zero and variance v. . Exercise 11.2.[2, p.189] Show that the mutual information I(X; Y ), in the case of this optimized distribution, is C=

 1 v log 1 + 2 . 2 σ

(11.26)

This is an important result. We see that the capacity of the Gaussian channel is a function of the signal-to-noise ratio v/σ 2 .

Inferences given a Gaussian input distribution If P (x) = Normal(x; 0, v) and P (y | x) = Normal(y; x, σ 2 ) then the marginal distribution of y is P (y) = Normal(y; 0, v+σ 2 ) and the posterior distribution of the input, given that the output is y, is: P (x | y) ∝ P (y | x)P (x)

(11.27) 2

2

2

∝ exp(−(y − x) /2σ ) exp(−x /2v)  !  v 1 1 −1 . = Normal x; y, + v + σ2 v σ2

(11.28) (11.29)

[The step from (11.28) to (11.29) is made by completing the square in the exponent.] This formula deserves careful study. The mean of the posterior v distribution, v+σ 2 y, can be viewed as a weighted combination of the value that best fits the output, x = y, and the value that best fits the prior, x = 0: v 1/σ 2 1/v y= y+ 0. 2 2 v+σ 1/v + 1/σ 1/v + 1/σ 2

(11.30)

The weights 1/σ 2 and 1/v are the precisions of the two Gaussians that we multiplied together in equation (11.28): the prior and the likelihood. The precision of the posterior distribution is the sum of these two precisions. This is a general property: whenever two independent sources contribute information, via Gaussian distributions, about an unknown variable, the precisions add. [This is the dual to the better-known relationship ‘when independent variables are added, their variances add’.]

Noisy-channel coding theorem for the Gaussian channel We have evaluated a maximal mutual information. Does it correspond to a maximum possible rate of error-free information transmission? One way of proving that this is so is to define a sequence of discrete channels, all derived from the Gaussian channel, with increasing numbers of inputs and outputs, and prove that the maximum mutual information of these channels tends to the asserted C. The noisy-channel coding theorem for discrete channels applies to each of these derived channels, thus we obtain a coding theorem for the continuous channel. Alternatively, we can make an intuitive argument for the coding theorem specific for the Gaussian channel.

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182

11 — Error-Correcting Codes and Real Channels

Geometrical view of the noisy-channel coding theorem: sphere packing Consider a sequence x = (x1 , . . . , xN ) of inputs, and the corresponding output y, as defining two points in an N dimensional space. For large N , the noise power is very likely to be close (fractionally) to N σ 2 . The output y is therefore √ very likely to be close to the surface of a sphere of radius N σ 2 centred on x. Similarly, if the original signal x is generated at random subject to an average power constraint x2 = v, √ then x is likely to lie close to a sphere, centred on the origin, of radius N v; and because the total average power of y is v + σ 2 , the received signal y is likely to lie on the surface of a sphere of radius p N (v + σ 2 ), centred on the origin. The volume of an N -dimensional sphere of radius r is V (r, N ) =

π N/2 rN . Γ(N/2+1)

(11.31)

Now consider making a communication system based on non-confusable inputs x, that is, inputs whose spheres do not overlap significantly. The maximum number S of non-confusable inputs is given by dividing the volume of the sphere of probable ys by the volume of the sphere for y given x: !N p N (v + σ 2 ) √ N σ2

(11.32)

 1 v  1 log M ≤ log 1 + 2 . N 2 σ

(11.33)

S≤ Thus the capacity is bounded by: C=

A more detailed argument like the one used in the previous chapter can establish equality.

Back to the continuous channel

This formula gives insight into the tradeoffs of practical communication. Imagine that we have a fixed power constraint. What is the best bandwidth to make use of that power? Introducing W0 = P/N0 , i.e., the bandwidth for which the signal-to-noise ratio is 1, figure 11.5 shows C/W 0 = W/W0 log(1 + W0 /W ) as a function of W/W0 . The capacity increases to an asymptote of W 0 log e. It is dramatically better (in terms of capacity for fixed power) to transmit at a low signal-to-noise ratio over a large bandwidth, than with high signal-to-noise in a narrow bandwidth; this is one motivation for wideband communication methods such as the ‘direct sequence spread-spectrum’ approach used in 3G mobile phones. Of course, you are not alone, and your electromagnetic neighbours may not be pleased if you use a large bandwidth, so for social reasons, engineers often have to make do with higher-power, narrow-bandwidth transmitters.

capacity

Recall that the use of a real continuous channel with bandwidth W , noise spectral density N0 and power P is equivalent to N/T = 2W uses per second of a Gaussian channel with σ 2 = N0 /2 and subject to the constraint x2n ≤ P/2W . Substituting the result for the capacity of the Gaussian channel, we find the capacity of the continuous channel to be:   P bits per second. (11.34) C = W log 1 + N0 W

1.4 1.2 1 0.8 0.6 0.4 0.2 0 0

1

2

3

4

5

6

bandwidth

Figure 11.5. Capacity versus bandwidth for a real channel: C/W0 = W/W0 log (1 + W0 /W ) as a function of W/W0 .

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11.4: What are the capabilities of practical error-correcting codes?

183

11.4 What are the capabilities of practical error-correcting codes? Nearly all codes are good, but nearly all codes require exponential look-up tables for practical implementation of the encoder and decoder – exponential in the blocklength N . And the coding theorem required N to be large. By a practical error-correcting code, we mean one that can be encoded and decoded in a reasonable amount of time, for example, a time that scales as a polynomial function of the blocklength N – preferably linearly.

The Shannon limit is not achieved in practice The non-constructive proof of the noisy-channel coding theorem showed that good block codes exist for any noisy channel, and indeed that nearly all block codes are good. But writing down an explicit and practical encoder and decoder that are as good as promised by Shannon is still an unsolved problem. Very good codes. Given a channel, a family of block codes that achieve arbitrarily small probability of error at any communication rate up to the capacity of the channel are called ‘very good’ codes for that channel. Good codes are code families that achieve arbitrarily small probability of error at non-zero communication rates up to some maximum rate that may be less than the capacity of the given channel. Bad codes are code families that cannot achieve arbitrarily small probability of error, or that can achieve arbitrarily small probability of error only by decreasing the information rate to zero. Repetition codes are an example of a bad code family. (Bad codes are not necessarily useless for practical purposes.) Practical codes are code families that can be encoded and decoded in time and space polynomial in the blocklength.

Most established codes are linear codes Let us review the definition of a block code, and then add the definition of a linear block code. An (N, K) block code for a channel Q is a list of S = 2 K codewords K {x(1) , x(2) , . . . , x(2 ) }, each of length N : x(s) ∈ AN X . The signal to be encoded, s, which comes from an alphabet of size 2 K , is encoded as x(s) . A linear (N, K) block code is a block code in which the codewords {x (s) } make up a K-dimensional subspace of A N X . The encoding operation can be represented by an N × K binary matrix GT such that if the signal to be encoded, in binary notation, is s (a vector of length K bits), then the encoded signal is t = GTs modulo 2. The codewords {t} can be defined as the set of vectors satisfying Ht = 0 mod 2, where H is the parity-check matrix of the code. For example the (7, 4) Hamming code of section 1.2 takes K = 4 signal bits, s, and transmits them followed by three parity-check bits. The N = 7 transmitted symbols are given by GTs mod 2. Coding theory was born with the work of Hamming, who invented a family of practical error-correcting codes, each able to correct one error in a block of length N , of which the repetition code R 3 and the (7, 4) code are





1· · ·  ·· 1· 1· ··    GT =  · · · 1  111 ·  · 111 1 · 11

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184

11 — Error-Correcting Codes and Real Channels

the simplest. Since then most established codes have been generalizations of Hamming’s codes: Bose–Chaudhury–Hocquenhem codes, Reed–M¨ uller codes, Reed–Solomon codes, and Goppa codes, to name a few.

Convolutional codes Another family of linear codes are convolutional codes, which do not divide the source stream into blocks, but instead read and transmit bits continuously. The transmitted bits are a linear function of the past source bits. Usually the rule for generating the transmitted bits involves feeding the present source bit into a linear-feedback shift-register of length k, and transmitting one or more linear functions of the state of the shift register at each iteration. The resulting transmitted bit stream is the convolution of the source stream with a linear filter. The impulse-response function of this filter may have finite or infinite duration, depending on the choice of feedback shift-register. We will discuss convolutional codes in Chapter 48.

Are linear codes ‘good’ ? One might ask, is the reason that the Shannon limit is not achieved in practice because linear codes are inherently not as good as random codes? The answer is no, the noisy-channel coding theorem can still be proved for linear codes, at least for some channels (see Chapter 14), though the proofs, like Shannon’s proof for random codes, are non-constructive. Linear codes are easy to implement at the encoding end. Is decoding a linear code also easy? Not necessarily. The general decoding problem (find the maximum likelihood s in the equation GTs + n = r) is in fact NP-complete (Berlekamp et al., 1978). [NP-complete problems are computational problems that are all equally difficult and which are widely believed to require exponential computer time to solve in general.] So attention focuses on families of codes for which there is a fast decoding algorithm.

Concatenation One trick for building codes with practical decoders is the idea of concatenation. An encoder–channel–decoder system C → Q → D can be viewed as defining a super-channel Q0 with a smaller probability of error, and with complex correlations among its errors. We can create an encoder C 0 and decoder D 0 for this super-channel Q0 . The code consisting of the outer code C 0 followed by the inner code C is known as a concatenated code. Some concatenated codes make use of the idea of interleaving. We read the data in blocks, the size of each block being larger than the blocklengths of the constituent codes C and C 0 . After encoding the data of one block using code C 0 , the bits are reordered within the block in such a way that nearby bits are separated from each other once the block is fed to the second code C. A simple example of an interleaver is a rectangular code or product code in which the data are arranged in a K2 × K1 block, and encoded horizontally using an (N1 , K1 ) linear code, then vertically using a (N 2 , K2 ) linear code. . Exercise 11.3.[3 ] Show that either of the two codes can be viewed as the inner code or the outer code. As an example, figure 11.6 shows a product code in which we encode first with the repetition code R3 (also known as the Hamming code H(3, 1))

C 0 → C → Q → D → D0 {z } | Q0

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11.4: What are the capabilities of practical error-correcting codes? 1 0 1 1 0 0 (a) 1

1 0 1 1 0 0 1

1 0 1 1 0 0 1

? ? ? ? ? (b)

1 1 1 1 0 1 (c) 1

1 1 1 0 0 0 1

1 0 1 1 1 0 1

1 1 1 1 0 0 (d) 1 1 1 1 1 1 1 (d0 ) 1

1 1 1 1 0 0 1 0 1 1 0 0 0 1

1 1 1 1 0 0 1 1 0 1 1 0 0 1

1 1 1 0 0 0 1 1 1 1 1 1 0 0 0 0 0 0 (e) 1 1 1 1 1 1 (1)(1)(1) 1 1 1 1 1 1 0 0 0 0 0 0 (e0 ) 1 1 1

horizontally then with H(7, 4) vertically. The blocklength of the concatenated code is 27. The number of source bits per codeword is four, shown by the small rectangle. We can decode conveniently (though not optimally) by using the individual decoders for each of the subcodes in some sequence. It makes most sense to first decode the code which has the lowest rate and hence the greatest errorcorrecting ability. Figure 11.6(c–e) shows what happens if we receive the codeword of figure 11.6a with some errors (five bits flipped, as shown) and apply the decoder for H(3, 1) first, and then the decoder for H(7, 4). The first decoder corrects three of the errors, but erroneously modifies the third bit in the second row where there are two bit errors. The (7, 4) decoder can then correct all three of these errors. Figure 11.6(d0 – e0 ) shows what happens if we decode the two codes in the other order. In columns one and two there are two errors, so the (7, 4) decoder introduces two extra errors. It corrects the one error in column 3. The (3, 1) decoder then cleans up four of the errors, but erroneously infers the second bit.

Interleaving The motivation for interleaving is that by spreading out bits that are nearby in one code, we make it possible to ignore the complex correlations among the errors that are produced by the inner code. Maybe the inner code will mess up an entire codeword; but that codeword is spread out one bit at a time over several codewords of the outer code. So we can treat the errors introduced by the inner code as if they are independent.

Other channel models In addition to the binary symmetric channel and the Gaussian channel, coding theorists keep more complex channels in mind also. Burst-error channels are important models in practice. Reed–Solomon codes use Galois fields (see Appendix C.1) with large numbers of elements (e.g. 216 ) as their input alphabets, and thereby automatically achieve a degree of burst-error tolerance in that even if 17 successive bits are corrupted, only 2 successive symbols in the Galois field representation are corrupted. Concatenation and interleaving can give further protection against burst errors. The concatenated Reed–Solomon codes used on digital compact discs are able to correct bursts of errors of length 4000 bits.

185 Figure 11.6. A product code. (a) A string 1011 encoded using a concatenated code consisting of two Hamming codes, H(3, 1) and H(7, 4). (b) a noise pattern that flips 5 bits. (c) The received vector. (d) After decoding using the horizontal (3, 1) decoder, and (e) after subsequently using the vertical (7, 4) decoder. The decoded vector matches the original. (d0 , e0 ) After decoding in the other order, three errors still remain.

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186

11 — Error-Correcting Codes and Real Channels

. Exercise 11.4.[2, p.189] The technique of interleaving, which allows bursts of errors to be treated as independent, is widely used, but is theoretically a poor way to protect data against burst errors, in terms of the amount of redundancy required. Explain why interleaving is a poor method, using the following burst-error channel as an example. Time is divided into chunks of length N = 100 clock cycles; during each chunk, there is a burst with probability b = 0.2; during a burst, the channel is a binary symmetric channel with f = 0.5. If there is no burst, the channel is an error-free binary channel. Compute the capacity of this channel and compare it with the maximum communication rate that could conceivably be achieved if one used interleaving and treated the errors as independent. Fading channels are real channels like Gaussian channels except that the received power is assumed to vary with time. A moving mobile phone is an important example. The incoming radio signal is reflected off nearby objects so that there are interference patterns and the intensity of the signal received by the phone varies with its location. The received power can easily vary by 10 decibels (a factor of ten) as the phone’s antenna moves through a distance similar to the wavelength of the radio signal (a few centimetres).

11.5 The state of the art What are the best known codes for communicating over Gaussian channels? All the practical codes are linear codes, and are either based on convolutional codes or block codes.

Convolutional codes, and codes based on them Textbook convolutional codes. The ‘de facto standard’ error-correcting code for satellite communications is a convolutional code with constraint length 7. Convolutional codes are discussed in Chapter 48. Concatenated convolutional codes. The above convolutional code can be used as the inner code of a concatenated code whose outer code is a Reed– Solomon code with eight-bit symbols. This code was used in deep space communication systems such as the Voyager spacecraft. For further reading about Reed–Solomon codes, see Lin and Costello (1983). The code for Galileo. A code using the same format but using a longer constraint length – 15 – for its convolutional code and a larger Reed– Solomon code was developed by the Jet Propulsion Laboratory (Swanson, 1988). The details of this code are unpublished outside JPL, and the decoding is only possible using a room full of special-purpose hardware. In 1992, this was the best code known of rate 1/4. Turbo codes. In 1993, Berrou, Glavieux and Thitimajshima reported work on turbo codes. The encoder of a turbo code is based on the encoders of two convolutional codes. The source bits are fed into each encoder, the order of the source bits being permuted in a random way, and the resulting parity bits from each constituent code are transmitted. The decoding algorithm involves iteratively decoding each constituent code using its standard decoding algorithm, then using the output of the decoder as the input to the other decoder. This decoding algorithm

- C1 -

 - π C  2

Figure 11.7. The encoder of a turbo code. Each box C1 , C2 , contains a convolutional code. The source bits are reordered using a permutation π before they are fed to C2 . The transmitted codeword is obtained by concatenating or interleaving the outputs of the two convolutional codes. The random permutation is chosen when the code is designed, and fixed thereafter.

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11.6: Summary

187

is an instance of a message-passing algorithm called the sum–product algorithm. Turbo codes are discussed in Chapter 48, and message passing in Chapters 16, 17, 25, and 26.

H =

Block codes Gallager’s low-density parity-check codes. The best block codes known for Gaussian channels were invented by Gallager in 1962 but were promptly forgotten by most of the coding theory community. They were rediscovered in 1995 and shown to have outstanding theoretical and practical properties. Like turbo codes, they are decoded by message-passing algorithms. We will discuss these beautifully simple codes in Chapter 47. The performances of the above codes are compared for Gaussian channels in figure 47.17, p.568.

11.6 Summary Random codes are good, but they require exponential resources to encode and decode them. Non-random codes tend for the most part not to be as good as random codes. For a non-random code, encoding may be easy, but even for simply-defined linear codes, the decoding problem remains very difficult. The best practical codes (a) employ very large block sizes; (b) are based on semi-random code constructions; and (c) make use of probabilitybased decoding algorithms.

11.7 Nonlinear codes Most practically used codes are linear, but not all. Digital soundtracks are encoded onto cinema film as a binary pattern. The likely errors affecting the film involve dirt and scratches, which produce large numbers of 1s and 0s respectively. We want none of the codewords to look like all-1s or all-0s, so that it will be easy to detect errors caused by dirt and scratches. One of the codes used in digital cinema sound systems is a nonlinear (8, 6) code consisting  of 64 of the 84 binary patterns of weight 4.

11.8 Errors other than noise

Another source of uncertainty for the receiver is uncertainty about the timing of the transmitted signal x(t). In ordinary coding theory and information theory, the transmitter’s time t and the receiver’s time u are assumed to be perfectly synchronized. But if the receiver receives a signal y(u), where the receiver’s time, u, is an imperfectly known function u(t) of the transmitter’s time t, then the capacity of this channel for communication is reduced. The theory of such channels is incomplete, compared with the synchronized channels we have discussed thus far. Not even the capacity of channels with synchronization errors is known (Levenshtein, 1966; Ferreira et al., 1997); codes for reliable communication over channels with synchronization errors remain an active research area (Davey and MacKay, 2001).

Figure 11.8. A low-density parity-check matrix and the corresponding graph of a rate-1/4 low-density parity-check code with blocklength N = 16, and M = 12 constraints. Each white circle represents a transmitted bit. Each bit participates in j = 3 constraints, represented by squares. Each constraint forces the sum of the k = 4 bits to which it is connected to be even. This code is a (16, 4) code. Outstanding performance is obtained when the blocklength is increased to N ' 10 000.

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188

11 — Error-Correcting Codes and Real Channels

Further reading For a review of the history of spread-spectrum methods, see Scholtz (1982).

11.9 Exercises The Gaussian channel . Exercise 11.5.[2, p.190] Consider a Gaussian channel with a real input x, and signal to noise ratio v/σ 2 . (a) What is its capacity C?

√ (b) If the input is constrained to be binary, x ∈ {± v}, what is the capacity C 0 of this constrained channel? (c) If in addition the output of the channel is thresholded using the mapping  1 y>0 0 y→y = (11.35) 0 y ≤ 0, what is the capacity C 00 of the resulting channel? (d) Plot the three capacities above as a function of v/σ 2 from 0.1 to 2. [You’ll need to do a numerical integral to evaluate C 0 .] . Exercise 11.6.[3 ] For large integers K and N , what fraction of all binary errorcorrecting codes of length N and rate R = K/N are linear codes? [The answer will depend on whether you choose to define the code to be an ordered list of 2K codewords, that is, a mapping from s ∈ {1, 2, . . . , 2 K } to x(s) , or to define the code to be an unordered list, so that two codes consisting of the same codewords are identical. Use the latter definition: a code is a set of codewords; how the encoder operates is not part of the definition of the code.]

Erasure channels . Exercise 11.7.[4 ] Design a code for the binary erasure channel, and a decoding algorithm, and evaluate their probability of error. [The design of good codes for erasure channels is an active research area (Spielman, 1996; Byers et al., 1998); see also Chapter 50.] . Exercise 11.8.[5 ] Design a code for the q-ary erasure channel, whose input x is drawn from 0, 1, 2, 3, . . . , (q − 1), and whose output y is equal to x with probability (1 − f ) and equal to ? otherwise. [This erasure channel is a good model for packets transmitted over the internet, which are either received reliably or are lost.] Exercise 11.9.[3, p.190] How do redundant arrays of independent disks (RAID) work? These are information storage systems consisting of about ten disk drives, of which any two or three can be disabled and the others are able to still able to reconstruct any requested file. What codes are used, and how far are these systems from the Shannon limit for the problem they are solving? How would you design a better RAID system? Some information is provided in the solution section. See http://www.acnc. com/raid2.html; see also Chapter 50.

[Some people say RAID stands for ‘redundant array of inexpensive disks’, but I think that’s silly – RAID would still be a good idea even if the disks were expensive!]

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11.10: Solutions

189

11.10 Solutions Solution to exercise 11.1 (p.181). Introduce a Lagrange multiplier λ for the power constraint and another, µ, for the constraint of normalization of P (x). R R F = I(X; Y ) − λ dx P (x)x2 − µ dx P (x) (11.36)  Z Z P (y | x) dy P (y | x) ln − λx2 − µ . (11.37) = dx P (x) P (y) Make the functional derivative with respect to P (x ∗ ). Z δF P (y | x∗ ) = dy P (y | x∗ ) ln − λx∗2 − µ ∗ δP (x ) P (y) Z Z 1 δP (y) . (11.38) − dx P (x) dy P (y | x) P (y) δP (x∗ ) R The final factor δP (y)/δP (x∗ ) is found, using P (y) = dx P (x)P (y | x), to be P (y | x∗ ), and the whole of the last term collapses in a puff of smoke to 1, which can be absorbed into the µ term. √ Substitute P (y | x) = exp(−(y − x)2 /2σ 2 )/ 2πσ 2 and set the derivative to zero: Z P (y | x) − λx2 − µ0 = 0 (11.39) dy P (y | x) ln P (y) Z 1 exp(−(y − x)2 /2σ 2 ) √ ln [P (y)σ] = −λx2 − µ0 − . (11.40) ⇒ dy 2 2 2πσ

This condition must be satisfied by ln[P (y)σ] for all x. Writing a Taylor expansion of ln[P (y)σ] = a+by+cy 2 +· · ·, only a quadratic function ln[P (y)σ] = a + cy 2 would satisfy the constraint (11.40). (Any higher order terms y p , p > 2, would produce terms in xp that are not present on the right-hand side.) Therefore P (y) is Gaussian. We can obtain this optimal output distribution by using a Gaussian input distribution P (x). Solution to exercise 11.2 (p.181). Given a Gaussian input distribution of variance v, the output distribution is Normal(0, v + σ 2 ), since x and the noise are independent random variables, and variances add for independent random variables. The mutual information is: Z Z I(X; Y ) = dx dy P (x)P (y | x) log P (y | x) − dy P (y) log P (y) (11.41) = =

1 1 1 1 log 2 − log 2 σ 2 v + σ2  v  1 log 1 + 2 . 2 σ

(11.42) (11.43)

Solution to exercise 11.4 (p.186). The capacity of the channel is one minus the information content of the noise that it adds. That information content is, per chunk, the entropy of the selection of whether the chunk is bursty, H 2 (b), plus, with probability b, the entropy of the flipped bits, N , which adds up to H2 (b) + N b per chunk (roughly; accurate if N is large). So, per bit, the capacity is, for N = 100,   1 C =1− H2 (b) + b = 1 − 0.207 = 0.793. (11.44) N In contrast, interleaving, which treats bursts of errors as independent, causes the channel to be treated as a binary symmetric channel with f = 0.2 × 0.5 = 0.1, whose capacity is about 0.53.

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11 — Error-Correcting Codes and Real Channels

Interleaving throws away the useful information about the correlatedness of the errors. Theoretically, we should be able to communicate about (0.79/0.53) ' 1.6 times faster using a code and decoder that explicitly treat bursts as bursts. Solution to exercise 11.5 (p.188). (a) Putting together the results of exercises 11.1 and 11.2, we deduce that a Gaussian channel with real input x, and signal to noise ratio v/σ 2 has capacity  v 1 (11.45) C = log 1 + 2 . 2 σ √ (b) If the input is constrained to be binary, x ∈ {± v}, the capacity is achieved by using these two inputs with equal probability. The capacity is reduced to a somewhat messy integral, Z ∞ Z ∞ dy P (y) log P (y), (11.46) dy N (y; 0) log N (y; 0) − C 00 = −∞



−∞

√ where N (y; x) ≡ (1/ 2π) exp[(y − x)2 /2], x ≡ v/σ, and P (y) ≡ [N (y; x) + N (y; −x)]/2. This capacity is smaller than the unconstrained capacity (11.45), but for small signal-to-noise ratio, the two capacities are close in value.

(c) If the output is thresholded, then the Gaussian channel is turned into a binary symmetric channel whose transition probability is given by the error function Φ defined on page 156. The capacity is √ C 00 = 1 − H2 (f ), where f = Φ( v/σ).

(11.47)

Solution to exercise 11.9 (p.188). There are several RAID systems. One of the easiest to understand consists of 7 disk drives which store data at rate 4/7 using a (7, 4) Hamming code: each successive four bits are encoded with the code and the seven codeword bits are written one to each disk. Two or perhaps three disk drives can go down and the others can recover the data. The effective channel model here is a binary erasure channel, because it is assumed that we can tell when a disk is dead. It is not possible to recover the data for some choices of the three dead disk drives; can you see why? . Exercise 11.10. [2, p.190] Give an example of three disk drives that, if lost, lead to failure of the above RAID system, and three that can be lost without failure. Solution to exercise 11.10 (p.190). The (7, 4) Hamming code has codewords of weight 3. If any set of three disk drives corresponding to one of those codewords is lost, then the other four disks can recover only 3 bits of information about the four source bits; a fourth bit is lost. [cf. exercise 13.13 (p.220) with q = 2: there are no binary MDS codes. This deficit is discussed further in section 13.11.] Any other set of three disk drives can be lost without problems because the corresponding four by four submatrix of the generator matrix is invertible. A better code would be a digital fountain – see Chapter 50.

1.2 1 0.8 0.6 0.4 0.2 0 0

0.5

1

1.5

2

2.5

1

0.1

0.01 0.1

1

Figure 11.9. Capacities (from top to bottom in each graph) C, C 0 , and C 00√, versus the signal-to-noise ratio ( v/σ). The lower graph is a log–log plot.

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Part III

Further Topics in Information Theory

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About Chapter 12 In Chapters 1–11, we concentrated on two aspects of information theory and coding theory: source coding – the compression of information so as to make efficient use of data transmission and storage channels; and channel coding – the redundant encoding of information so as to be able to detect and correct communication errors. In both these areas we started by ignoring practical considerations, concentrating on the question of the theoretical limitations and possibilities of coding. We then discussed practical source-coding and channel-coding schemes, shifting the emphasis towards computational feasibility. But the prime criterion for comparing encoding schemes remained the efficiency of the code in terms of the channel resources it required: the best source codes were those that achieved the greatest compression; the best channel codes were those that communicated at the highest rate with a given probability of error. In this chapter we now shift our viewpoint a little, thinking of ease of information retrieval as a primary goal. It turns out that the random codes which were theoretically useful in our study of channel coding are also useful for rapid information retrieval. Efficient information retrieval is one of the problems that brains seem to solve effortlessly, and content-addressable memory is one of the topics we will study when we look at neural networks.

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12 Hash Codes: Codes for Efficient Information Retrieval 12.1 The information-retrieval problem A simple example of an information-retrieval problem is the task of implementing a phone directory service, which, in response to a person’s name, returns (a) a confirmation that that person is listed in the directory; and (b) the person’s phone number and other details. We could formalize this problem as follows, with S being the number of names that must be stored in the directory. You are given a list of S binary strings of length N bits, {x (1) , . . . , x(S) }, where S is considerably smaller than the total number of possible strings, 2 N . We will call the superscript ‘s’ in x (s) the record number of the string. The idea is that s runs over customers in the order in which they are added to the directory and x(s) is the name of customer s. We assume for simplicity that all people have names of the same length. The name length might be, say, N = 200 bits, and we might want to store the details of ten million customers, so S ' 107 ' 223 . We will ignore the possibility that two customers have identical names. The task is to construct the inverse of the mapping from s to x (s) , i.e., to make a system that, given a string x, returns the value of s such that x = x (s) if one exists, and otherwise reports that no such s exists. (Once we have the record number, we can go and look in memory location s in a separate memory full of phone numbers to find the required number.) The aim, when solving this task, is to use minimal computational resources in terms of the amount of memory used to store the inverse mapping from x to s and the amount of time to compute the inverse mapping. And, preferably, the inverse mapping should be implemented in such a way that further new strings can be added to the directory in a small amount of computer time too.

Some standard solutions The simplest and dumbest solutions to the information-retrieval problem are a look-up table and a raw list. The look-up table is a piece of memory of size 2 N log2 S, log 2 S being the amount of memory required to store an integer between 1 and S. In each of the 2N locations, we put a zero, except for the locations x that correspond to strings x(s) , into which we write the value of s. The look-up table is a simple and quick solution, but only if there is sufficient memory for the table, and if the cost of looking up entries in 193

string length number of strings number of possible strings

N ' 200 S ' 223 2N ' 2200

Figure 12.1. Cast of characters.

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12 — Hash Codes: Codes for Efficient Information Retrieval memory is independent of the memory size. But in our definition of the task, we assumed that N is about 200 bits or more, so the amount of memory required would be of size 2200 ; this solution is completely out of the question. Bear in mind that the number of particles in the solar system is only about 2190 .

The raw list is a simple list of ordered pairs (s, x (s) ) ordered by the value of s. The mapping from x to s is achieved by searching through the list of strings, starting from the top, and comparing the incoming string x with each record x(s) until a match is found. This system is very easy to maintain, and uses a small amount of memory, about SN bits, but is rather slow to use, since on average five million pairwise comparisons will be made. . Exercise 12.1.[2, p.202] Show that the average time taken to find the required string in a raw list, assuming that the original names were chosen at random, is about S + N binary comparisons. (Note that you don’t have to compare the whole string of length N , since a comparison can be terminated as soon as a mismatch occurs; show that you need on average two binary comparisons per incorrect string match.) Compare this with the worst-case search time – assuming that the devil chooses the set of strings and the search key. The standard way in which phone directories are made improves on the look-up table and the raw list by using an alphabetically-ordered list. Alphabetical list. The strings {x(s) } are sorted into alphabetical order. Searching for an entry now usually takes less time than was needed for the raw list because we can take advantage of the sortedness; for example, we can open the phonebook at its middle page, and compare the name we find there with the target string; if the target is ‘greater’ than the middle string then we know that the required string, if it exists, will be found in the second half of the alphabetical directory. Otherwise, we look in the first half. By iterating this splitting-in-the-middle procedure, we can identify the target string, or establish that the string is not listed, in dlog 2 Se string comparisons. The expected number of binary comparisons per string comparison will tend to increase as the search progresses, but the total number of binary comparisons required will be no greater than dlog 2 SeN . The amount of memory required is the same as that required for the raw list.

Adding new strings to the database requires that we insert them in the correct location in the list. To find that location takes about dlog 2 Se binary comparisons. Can we improve on the well-established alphabetized list? Let us consider our task from some new viewpoints. The task is to construct a mapping x → s from N bits to log 2 S bits. This is a pseudo-invertible mapping, since for any x that maps to a non-zero s, the customer database contains the pair (s, x (s) ) that takes us back. Where have we come across the idea of mapping from N bits to M bits before? We encountered this idea twice: first, in source coding, we studied block codes which were mappings from strings of N symbols to a selection of one label in a list. The task of information retrieval is similar to the task (which

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12.2: Hash codes

195

we never actually solved) of making an encoder for a typical-set compression code. The second time that we mapped bit strings to bit strings of another dimensionality was when we studied channel codes. There, we considered codes that mapped from K bits to N bits, with N greater than K, and we made theoretical progress using random codes. In hash codes, we put together these two notions. We will study random codes that map from N bits to M bits where M is smaller than N . The idea is that we will map the original high-dimensional space down into a lower-dimensional space, one in which it is feasible to implement the dumb look-up table method which we rejected a moment ago.

12.2 Hash codes First we will describe how a hash code works, then we will study the properties of idealized hash codes. A hash code implements a solution to the informationretrieval problem, that is, a mapping from x to s, with the help of a pseudorandom function called a hash function, which maps the N -bit string x to an M -bit string h(x), where M is smaller than N . M is typically chosen such that the ‘table size’ T ' 2M is a little bigger than S – say, ten times bigger. For example, if we were expecting S to be about a million, we might map x into a 30-bit hash h (regardless of the size N of each item x). The hash function is some fixed deterministic function which should ideally be indistinguishable from a fixed random code. For practical purposes, the hash function must be quick to compute. Two simple examples of hash functions are: Division method. The table size T is a prime number, preferably one that is not close to a power of 2. The hash value is the remainder when the integer x is divided by T . Variable string addition method. This method assumes that x is a string of bytes and that the table size T is 256. The characters of x are added, modulo 256. This hash function has the defect that it maps strings that are anagrams of each other onto the same hash. It may be improved by putting the running total through a fixed pseudorandom permutation after each character is added. In the variable string exclusive-or method with table size ≤ 65 536, the string is hashed twice in this way, with the initial running total being set to 0 and 1 respectively (algorithm 12.3). The result is a 16-bit hash. Having picked a hash function h(x), we implement an information retriever as follows. (See figure 12.4.) Encoding. A piece of memory called the hash table is created of size 2 M b memory units, where b is the amount of memory needed to represent an integer between 0 and S. This table is initially set to zero throughout. Each memory x(s) is put through the hash function, and at the location in the hash table corresponding to the resulting vector h (s) = h(x(s) ), the integer s is written – unless that entry in the hash table is already occupied, in which case we have a collision between x (s) and some earlier 0 x(s ) which both happen to have the same hash code. Collisions can be handled in various ways – we will discuss some in a moment – but first let us complete the basic picture.

string length number of strings size of hash function size of hash table

N ' 200 S ' 223 M ' 30 bits T = 2M ' 230

Figure 12.2. Revised cast of characters.

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12 — Hash Codes: Codes for Efficient Information Retrieval

unsigned char Rand8[256]; int Hash(char *x) { int h; unsigned char h1, h2; if (*x == 0) return 0; h1 = *x; h2 = *x + 1; x++; while (*x) { h1 = Rand8[h1 ^ *x]; h2 = Rand8[h2 ^ *x]; x++; } h = ((int)(h1)<<8) | (int) h2 ; return h ;

// This array contains a random permutation from 0..255 to 0..255 // x is a pointer to the first char; // *x is the first character

Algorithm 12.3. C code implementing the variable string exclusive-or method to create a hash h in the range 0 . . . 65 535 from a string x. Author: Thomas Niemann.

// Special handling of empty string // Initialize two hashes // Proceed to the next character // Exclusive-or with the two hashes // and put through the randomizer // End of string is reached when *x=0 // Shift h1 left 8 bits and add h2 // Hash is concatenation of h1 and h2

}

Hash function

- hashes

Strings

 6

N bits x(1) x(2) x(3) .. .

S x

 

@

A A

2

h(x(1) ) →

1

h(x(3) ) →

3

A U A h(x(s) ) →

s

  

@ R @

A A

M bits

 -

h(x(2) ) →

 @  @ R @

(s)

.. .

?

-

Hash table

6

2M

?

Figure 12.4. Use of hash functions for information retrieval. For each string x(s) , the hash h = h(x(s) ) is computed, and the value of s is written into the hth row of the hash table. Blank rows in the hash table contain the value zero. The table size is T = 2M .

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12.3: Collision resolution Decoding. To retrieve a piece of information corresponding to a target vector x, we compute the hash h of x and look at the corresponding location in the hash table. If there is a zero, then we know immediately that the string x is not in the database. The cost of this answer is the cost of one hash-function evaluation and one look-up in the table of size 2 M . If, on the other hand, there is a non-zero entry s in the table, there are two possibilities: either the vector x is indeed equal to x (s) ; or the vector x(s) is another vector that happens to have the same hash code as the target x. (A third possibility is that this non-zero entry might have something to do with our yet-to-be-discussed collision-resolution system.) To check whether x is indeed equal to x (s) , we take the tentative answer s, look up x(s) in the original forward database, and compare it bit by bit with x; if it matches then we report s as the desired answer. This successful retrieval has an overall cost of one hash-function evaluation, one look-up in the table of size 2M , another look-up in a table of size S, and N binary comparisons – which may be much cheaper than the simple solutions presented in section 12.1. . Exercise 12.2.[2, p.202] If we have checked the first few bits of x (s) with x and found them to be equal, what is the probability that the correct entry has been retrieved, if the alternative hypothesis is that x is actually not in the database? Assume that the original source strings are random, and the hash function is a random hash function. How many binary evaluations are needed to be sure with odds of a billion to one that the correct entry has been retrieved? The hashing method of information retrieval can be used for strings x of arbitrary length, if the hash function h(x) can be applied to strings of any length.

12.3 Collision resolution We will study two ways of resolving collisions: appending in the table, and storing elsewhere.

Appending in table When encoding, if a collision occurs, we continue down the hash table and write the value of s into the next available location in memory that currently contains a zero. If we reach the bottom of the table before encountering a zero, we continue from the top. When decoding, if we compute the hash code for x and find that the s contained in the table doesn’t point to an x (s) that matches the cue x, we continue down the hash table until we either find an s whose x (s) does match the cue x, in which case we are done, or else encounter a zero, in which case we know that the cue x is not in the database. For this method, it is essential that the table be substantially bigger in size than S. If 2M < S then the encoding rule will become stuck with nowhere to put the last strings.

Storing elsewhere A more robust and flexible method is to use pointers to additional pieces of memory in which collided strings are stored. There are many ways of doing

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12 — Hash Codes: Codes for Efficient Information Retrieval

this. As an example, we could store in location h in the hash table a pointer (which must be distinguishable from a valid record number s) to a ‘bucket’ where all the strings that have hash code h are stored in a sorted list. The encoder sorts the strings in each bucket alphabetically as the hash table and buckets are created. The decoder simply has to go and look in the relevant bucket and then check the short list of strings that are there by a brief alphabetical search. This method of storing the strings in buckets allows the option of making the hash table quite small, which may have practical benefits. We may make it so small that almost all strings are involved in collisions, so all buckets contain a small number of strings. It only takes a small number of binary comparisons to identify which of the strings in the bucket matches the cue x.

12.4 Planning for collisions: a birthday problem Exercise 12.3.[2, p.202] If we wish to store S entries using a hash function whose output has M bits, how many collisions should we expect to happen, assuming that our hash function is an ideal random function? What size M of hash table is needed if we would like the expected number of collisions to be smaller than 1? What size M of hash table is needed if we would like the expected number of collisions to be a small fraction, say 1%, of S? [Notice the similarity of this problem to exercise 9.20 (p.156).]

12.5 Other roles for hash codes Checking arithmetic If you wish to check an addition that was done by hand, you may find useful the method of casting out nines. In casting out nines, one finds the sum, modulo nine, of all the digits of the numbers to be summed and compares it with the sum, modulo nine, of the digits of the putative answer. [With a little practice, these sums can be computed much more rapidly than the full original addition.] Example 12.4. In the calculation shown in the margin the sum, modulo nine, of the digits in 189+1254+238 is 7, and the sum, modulo nine, of 1+6+8+1 is 7. The calculation thus passes the casting-out-nines test. Casting out nines gives a simple example of a hash function. For any addition expression of the form a + b + c + · · ·, where a, b, c, . . . are decimal numbers we define h ∈ {0, 1, 2, 3, 4, 5, 6, 7, 8} by h(a + b + c + · · ·) = sum modulo nine of all digits in a, b, c ;

(12.1)

then it is nice property of decimal arithmetic that if a + b +c + ··· = m + n + o + ···

(12.2)

then the hashes h(a + b + c + · · ·) and h(m + n + o + · · ·) are equal. . Exercise 12.5.[1, p.203] What evidence does a correct casting-out-nines match give in favour of the hypothesis that the addition has been done correctly?

189 +1254 + 238 1681

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12.5: Other roles for hash codes

Error detection among friends Are two files the same? If the files are on the same computer, we could just compare them bit by bit. But if the two files are on separate machines, it would be nice to have a way of confirming that two files are identical without having to transfer one of the files from A to B. [And even if we did transfer one of the files, we would still like a way to confirm whether it has been received without modifications!] This problem can be solved using hash codes. Let Alice and Bob be the holders of the two files; Alice sent the file to Bob, and they wish to confirm it has been received without error. If Alice computes the hash of her file and sends it to Bob, and Bob computes the hash of his file, using the same M -bit hash function, and the two hashes match, then Bob can deduce that the two files are almost surely the same. Example 12.6. What is the probability of a false negative, i.e., the probability, given that the two files do differ, that the two hashes are nevertheless identical? If we assume that the hash function is random and that the process that causes the files to differ knows nothing about the hash function, then the probability of a false negative is 2−M . 2 A 32-bit hash gives a probability of false negative of about 10 −10 . It is common practice to use a linear hash function called a 32-bit cyclic redundancy check to detect errors in files. (A cyclic redundancy check is a set of 32 paritycheck bits similar to the 3 parity-check bits of the (7, 4) Hamming code.) To have a false-negative rate smaller than one in a billion, M = 32 bits is plenty, if the errors are produced by noise. . Exercise 12.7.[2, p.203] Such a simple parity-check code only detects errors; it doesn’t help correct them. Since error-correcting codes exist, why not use one of them to get some error-correcting capability too?

Tamper detection What if the differences between the two files are not simply ‘noise’, but are introduced by an adversary, a clever forger called Fiona, who modifies the original file to make a forgery that purports to be Alice’s file? How can Alice make a digital signature for the file so that Bob can confirm that no-one has tampered with the file? And how can we prevent Fiona from listening in on Alice’s signature and attaching it to other files? Let’s assume that Alice computes a hash function for the file and sends it securely to Bob. If Alice computes a simple hash function for the file like the linear cyclic redundancy check, and Fiona knows that this is the method of verifying the file’s integrity, Fiona can make her chosen modifications to the file and then easily identify (by linear algebra) a further 32-or-so single bits that, when flipped, restore the hash function of the file to its original value. Linear hash functions give no security against forgers. We must therefore require that the hash function be hard to invert so that no-one can construct a tampering that leaves the hash function unaffected. We would still like the hash function to be easy to compute, however, so that Bob doesn’t have to do hours of work to verify every file he received. Such a hash function – easy to compute, but hard to invert – is called a one-way

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12 — Hash Codes: Codes for Efficient Information Retrieval

hash function. Finding such functions is one of the active research areas of cryptography. A hash function that is widely used in the free software community to confirm that two files do not differ is MD5, which produces a 128-bit hash. The details of how it works are quite complicated, involving convoluted exclusiveor-ing and if-ing and and-ing.1 Even with a good one-way hash function, the digital signatures described above are still vulnerable to attack, if Fiona has access to the hash function. Fiona could take the tampered file and hunt for a further tiny modification to it such that its hash matches the original hash of Alice’s file. This would take some time – on average, about 232 attempts, if the hash function has 32 bits – but eventually Fiona would find a tampered file that matches the given hash. To be secure against forgery, digital signatures must either have enough bits for such a random search to take too long, or the hash function itself must be kept secret. Fiona has to hash 2M files to cheat. 232 file modifications is not very many, so a 32-bit hash function is not large enough for forgery prevention. Another person who might have a motivation for forgery is Alice herself. For example, she might be making a bet on the outcome of a race, without wishing to broadcast her prediction publicly; a method for placing bets would be for her to send to Bob the bookie the hash of her bet. Later on, she could send Bob the details of her bet. Everyone can confirm that her bet is consistent with the previously publicized hash. [This method of secret publication was used by Isaac Newton and Robert Hooke when they wished to establish priority for scientific ideas without revealing them. Hooke’s hash function was alphabetization as illustrated by the conversion of UT TENSIO, SIC VIS into the anagram CEIIINOSSSTTUV.] Such a protocol relies on the assumption that Alice cannot change her bet after the event without the hash coming out wrong. How big a hash function do we need to use to ensure that Alice cannot cheat? The answer is different from the size of the hash we needed in order to defeat Fiona above, because Alice is the author of both files. Alice could cheat by searching for two files that have identical hashes to each other. For example, if she’d like to cheat by placing two bets for the price of one, she could make a large number N1 of versions of bet one (differing from each other in minor details only), and a large number N 2 of versions of bet two, and hash them all. If there’s a collision between the hashes of two bets of different types, then she can submit the common hash and thus buy herself the option of placing either bet. Example 12.8. If the hash has M bits, how big do N 1 and N2 need to be for Alice to have a good chance of finding two different bets with the same hash? This is a birthday problem like exercise 9.20 (p.156). If there are N 1 Montagues and N2 Capulets at a party, and each is assigned a ‘birthday’ of M bits, the expected number of collisions between a Montague and a Capulet is N1 N2 2−M , 1

http://www.freesoft.org/CIE/RFC/1321/3.htm

(12.3)

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12.6: Further exercises

201

so to minimize the number of files hashed, N 1 + N2 , Alice should make N1 and N2 equal, and will need to hash about 2M/2 files until she finds two that match. 2 Alice has to hash 2M/2 files to cheat. [This is the square root of the number of hashes Fiona had to make.] If Alice has the use of C = 106 computers for T = 10 years, each computer taking t = 1 ns to evaluate a hash, the bet-communication system is secure against Alice’s dishonesty only if M  2 log 2 CT /t ' 160 bits.

Further reading The Bible for hash codes is volume 3 of Knuth (1968). I highly recommend the story of Doug McIlroy’s spell program, as told in section 13.8 of Programming Pearls (Bentley, 2000). This astonishing piece of software makes use of a 64kilobyte data structure to store the spellings of all the words of 75 000-word dictionary.

12.6 Further exercises Exercise 12.9.[1 ] What is the shortest the address on a typical international letter could be, if it is to get to a unique human recipient? (Assume the permitted characters are [A-Z,0-9].) How long are typical email addresses? Exercise 12.10. [2, p.203] How long does a piece of text need to be for you to be pretty sure that no human has written that string of characters before? How many notes are there in a new melody that has not been composed before? . Exercise 12.11. [3, p.204] Pattern recognition by molecules. Some proteins produced in a cell have a regulatory role. A regulatory protein controls the transcription of specific genes in the genome. This control often involves the protein’s binding to a particular DNA sequence in the vicinity of the regulated gene. The presence of the bound protein either promotes or inhibits transcription of the gene. (a) Use information-theoretic arguments to obtain a lower bound on the size of a typical protein that acts as a regulator specific to one gene in the whole human genome. Assume that the genome is a sequence of 3 × 109 nucleotides drawn from a four letter alphabet {A, C, G, T}; a protein is a sequence of amino acids drawn from a twenty letter alphabet. [Hint: establish how long the recognized DNA sequence has to be in order for that sequence to be unique to the vicinity of one gene, treating the rest of the genome as a random sequence. Then discuss how big the protein must be to recognize a sequence of that length uniquely.] (b) Some of the sequences recognized by DNA-binding regulatory proteins consist of a subsequence that is repeated twice or more, for example the sequence GCCCCCCACCCCTGCCCCC

(12.4)

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202

12 — Hash Codes: Codes for Efficient Information Retrieval is a binding site found upstream of the alpha-actin gene in humans. Does the fact that some binding sites consist of a repeated subsequence influence your answer to part (a)?

12.7 Solutions Solution to exercise 12.1 (p.194). First imagine comparing the string x with another random string x(s) . The probability that the first bits of the two strings match is 1/2. The probability that the second bits match is 1/2. Assuming we stop comparing once we hit the first mismatch, the expected number of matches is 1, so the expected number of comparisons is 2 (exercise 2.34, p.38). Assuming the correct string is located at random in the raw list, we will have to compare with an average of S/2 strings before we find it, which costs 2S/2 binary comparisons; and comparing the correct strings takes N binary comparisons, giving a total expectation of S + N binary comparisons, if the strings are chosen at random. In the worst case (which may indeed happen in practice), the other strings are very similar to the search key, so that a lengthy sequence of comparisons is needed to find each mismatch. The worst case is when the correct string is last in the list, and all the other strings differ in the last bit only, giving a requirement of SN binary comparisons. Solution to exercise 12.2 (p.197). The likelihood ratio for the two hypotheses, H0 : x(s) = x, and H1 : x(s) 6= x, contributed by the datum ‘the first bits of x(s) and x are equal’ is 1 P (Datum | H0 ) = = 2. P (Datum | H1 ) 1/2

(12.5)

If the first r bits all match, the likelihood ratio is 2 r to one. On finding that 30 bits match, the odds are a billion to one in favour of H 0 , assuming we start from even odds. [For a complete answer, we should compute the evidence given by the prior information that the hash entry s has been found in the table at h(x). This fact gives further evidence in favour of H 0 .] Solution to exercise 12.3 (p.198). Let the hash function have an output alphabet of size T = 2M . If M were equal to log 2 S then we would have exactly enough bits for each entry to have its own unique hash. The probability that one particular pair of entries collide under a random hash function is 1/T . The number of pairs is S(S − 1)/2. So the expected number of collisions between pairs is exactly S(S − 1)/(2T ). (12.6) If we would like this to be smaller than 1, then we need T > S(S − 1)/2 so M > 2 log 2 S.

(12.7)

We need twice as many bits as the number of bits, log 2 S, that would be sufficient to give each entry a unique name. If we are happy to have occasional collisions, involving a fraction f of the names S, then we need T > S/f (since the probability that one particular name is collided-with is f ' S/T ) so M > log 2 S + log 2 [1/f ],

(12.8)

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12.7: Solutions which means for f ' 0.01 that we need an extra 7 bits above log 2 S. The important point to note is the scaling of T with S in the two cases (12.7, 12.8). If we want the hash function to be collision-free, then we must have T greater than ∼ S 2 . If we are happy to have a small frequency of collisions, then T needs to be of order S only. Solution to exercise 12.5 (p.198). The posterior probability ratio for the two hypotheses, H+ = ‘calculation correct’ and H− = ‘calculation incorrect’ is the product of the prior probability ratio P (H + )/P (H− ) and the likelihood ratio, P (match | H+ )/P (match | H− ). This second factor is the answer to the question. The numerator P (match | H + ) is equal to 1. The denominator’s value depends on our model of errors. If we know that the human calculator is prone to errors involving multiplication of the answer by 10, or to transposition of adjacent digits, neither of which affects the hash value, then P (match | H − ) could be equal to 1 also, so that the correct match gives no evidence in favour of H+ . But if we assume that errors are ‘random from the point of view of the hash function’ then the probability of a false positive is P (match | H − ) = 1/9, and the correct match gives evidence 9:1 in favour of H + . Solution to exercise 12.7 (p.199). If you add a tiny M = 32 extra bits of hash to a huge N -bit file you get pretty good error detection – the probability that an error is undetected is 2−M , less than one in a billion. To do error correction requires far more check bits, the number depending on the expected types of corruption, and on the file size. For example, if just eight random bits in a 23  megabyte file are corrupted, it would take about log 2 28 ' 23 × 8 ' 180 bits to specify which are the corrupted bits, and the number of parity-check bits used by a successful error-correcting code would have to be at least this number, by the counting argument of exercise 1.10 (solution, p.20). Solution to exercise 12.10 (p.201). We want to know the length L of a string such that it is very improbable that that string matches any part of the entire writings of humanity. Let’s estimate that these writings total about one book for each person living, and that each book contains two million characters (200 pages with 10 000 characters per page) – that’s 10 16 characters, drawn from an alphabet of, say, 37 characters. The probability that a randomly chosen string of length L matches at one point in the collected works of humanity is 1/37 L . So the expected number of matches is 1016 /37L , which is vanishingly small if L ≥ 16/ log 10 37 ' 10. Because of the redundancy and repetition of humanity’s writings, it is possible that L ' 10 is an overestimate. So, if you want to write something unique, sit down and compose a string of ten characters. But don’t write gidnebinzz, because I already thought of that string. As for a new melody, if we focus on the sequence of notes, ignoring duration and stress, and allow leaps of up to an octave at each note, then the number of choices per note is 23. The pitch of the first note is arbitrary. The number of melodies of length r notes in this rather ugly ensemble of Sch¨onbergian tunes is 23r−1 ; for example, there are 250 000 of length r = 5. Restricting the permitted intervals will reduce this figure; including duration and stress will increase it again. [If we restrict the permitted intervals to repetitions and tones or semitones, the reduction is particularly severe; is this why the melody of ‘Ode to Joy’ sounds so boring?] The number of recorded compositions is probably less than a million. If you learn 100 new melodies per week for every week of your life then you will have learned 250 000 melodies at age 50. Based

203

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204

12 — Hash Codes: Codes for Efficient Information Retrieval

on empirical experience of playing the game ‘guess that tune’, it seems to me that whereas many four-note sequences are shared in common between melodies, the number of collisions between five-note sequences is rather smaller – most famous five-note sequences are unique.

In guess that tune, one player chooses a melody, and sings a gradually-increasing number of its notes, while the other participants try to guess the whole melody.

Solution to exercise 12.11 (p.201). (a) Let the DNA-binding protein recognize a sequence of length L nucleotides. That is, it binds preferentially to that DNA sequence, and not to any other pieces of DNA in the whole genome. (In reality, the recognized sequence may contain some wildcard characters, e.g., the * in TATAA*A, which denotes ‘any of A, C, G and T’; so, to be precise, we are assuming that the recognized sequence contains L non-wildcard characters.) Assuming the rest of the genome is ‘random’, i.e., that the sequence consists of random nucleotides A, C, G and T with equal probability – which is obviously untrue, but it shouldn’t make too much difference to our calculation – the chance that there is no other occurrence of the target sequence in the whole genome, of length N nucleotides, is roughly

The Parsons code is a related hash function for melodies: each pair of consecutive notes is coded as U (‘up’) if the second note is higher than the first, R (‘repeat’) if the pitches are equal, and D (‘down’) otherwise. You can find out how well this hash function works at http://musipedia.org/.

(1 − (1/4)L )N ' exp(−N (1/4)L ),

(12.9)

which is close to one only if that is,

N 4−L  1,

(12.10)

L > log N/ log 4.

(12.11)

Using N = 3 × 109 , we require the recognized sequence to be longer than Lmin = 16 nucleotides. What size of protein does this imply? • A weak lower bound can be obtained by assuming that the information content of the protein sequence itself is greater than the information content of the nucleotide sequence the protein prefers to bind to (which we have argued above must be at least 32 bits). This gives a minimum protein length of 32/ log 2 (20) ' 7 amino acids. • Thinking realistically, the recognition of the DNA sequence by the protein presumably involves the protein coming into contact with all sixteen nucleotides in the target sequence. If the protein is a monomer, it must be big enough that it can simultaneously make contact with sixteen nucleotides of DNA. One helical turn of DNA containing ten nucleotides has a length of 3.4 nm, so a contiguous sequence of sixteen nucleotides has a length of 5.4 nm. The diameter of the protein must therefore be about 5.4 nm or greater. Egg-white lysozyme is a small globular protein with a length of 129 amino acids and a diameter of about 4 nm. Assuming that volume is proportional to sequence length and that volume scales as the cube of the diameter, a protein of diameter 5.4 nm must have a sequence of length 2.5 × 129 ' 324 amino acids. (b) If, however, a target sequence consists of a twice-repeated sub-sequence, we can get by with a much smaller protein that recognizes only the sub-sequence, and that binds to the DNA strongly only if it can form a dimer, both halves of which are bound to the recognized sequence. Halving the diameter of the protein, we now only need a protein whose length is greater than 324/8 = 40 amino acids. A protein of length smaller than this cannot by itself serve as a regulatory protein specific to one gene, because it’s simply too small to be able to make a sufficiently specific match – its available surface does not have enough information content.

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About Chapter 13 In Chapters 8–11, we established Shannon’s noisy-channel coding theorem for a general channel with any input and output alphabets. A great deal of attention in coding theory focuses on the special case of channels with binary inputs. Constraining ourselves to these channels simplifies matters, and leads us into an exceptionally rich world, which we will only taste in this book. One of the aims of this chapter is to point out a contrast between Shannon’s aim of achieving reliable communication over a noisy channel and the apparent aim of many in the world of coding theory. Many coding theorists take as their fundamental problem the task of packing as many spheres as possible, with radius as large as possible, into an N -dimensional space, with no spheres overlapping. Prizes are awarded to people who find packings that squeeze in an extra few spheres. While this is a fascinating mathematical topic, we shall see that the aim of maximizing the distance between codewords in a code has only a tenuous relationship to Shannon’s aim of reliable communication.

205

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13 Binary Codes We’ve established Shannon’s noisy-channel coding theorem for a general channel with any input and output alphabets. A great deal of attention in coding theory focuses on the special case of channels with binary inputs, the first implicit choice being the binary symmetric channel. The optimal decoder for a code, given a binary symmetric channel, finds the codeword that is closest to the received vector, closest in Hamming distance. The Hamming distance between two binary vectors is the number of coordinates in which the two vectors differ. Decoding errors will occur if the noise takes us from the transmitted codeword t to a received vector r that is closer to some other codeword. The distances between codewords are thus relevant to the probability of a decoding error.

Example: The Hamming distance between 00001111 and 11001101 is 3.

13.1 Distance properties of a code The distance of a code is the smallest separation between two of its codewords. Example 13.1. The (7, 4) Hamming code (p.8) has distance d = 3. All pairs of its codewords differ in at least 3 bits. The maximum number of errors it can correct is t = 1; in general a code with distance d is b(d−1)/2cerror-correcting. A more precise term for distance is the minimum distance of the code. The distance of a code is often denoted by d or d min . We’ll now constrain our attention to linear codes. In a linear code, all codewords have identical distance properties, so we can summarize all the distances between the code’s codewords by counting the distances from the all-zero codeword. The weight enumerator function of a code, A(w), is defined to be the number of codewords in the code that have weight w. The weight enumerator function is also known as the distance distribution of the code.

w

A(w)

0 3 4 7

1 7 7 1

Total

16

8 7 6 5 4 3 2 1 0

Example 13.2. The weight enumerator functions of the (7, 4) Hamming code and the dodecahedron code are shown in figures 13.1 and 13.2.

13.2 Obsession with distance Since the maximum number of errors that a code can guarantee to correct, t, is related to its distance d by t = b(d−1)/2c, many coding theorists focus on the distance of a code, searching for codes of a given size that have the biggest possible distance. Much of practical coding theory has focused on decoders that give the optimal decoding for all error patterns of weight up to the half-distance t of their codes. 206

0

1

2

3

4

5

6

7

Figure 13.1. The graph of the (7, 4) Hamming code, and its weight enumerator function.

d = 2t + 1 if d is odd, and d = 2t + 2 if d is even.

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13.2: Obsession with distance

207 350

w

A(w)

300

0 5 8 9 10 11 12 13 14 15 16 17 18 19 20

1 12 30 20 72 120 100 180 240 272 345 300 200 120 36

250

Total

2048

200 150 100 50 0 0

5

8 10

15

20

25

30

0

5

8 10

15

20

25

30

100

Figure 13.2. The graph defining the (30, 11) dodecahedron code (the circles are the 30 transmitted bits and the triangles are the 20 parity checks, one of which is redundant) and the weight enumerator function (solid lines). The dotted lines show the average weight enumerator function of all random linear codes with the same size of generator matrix, which will be computed shortly. The lower figure shows the same functions on a log scale.

10

1

A bounded-distance decoder is a decoder that returns the closest codeword to a received binary vector r if the distance from r to that codeword is less than or equal to t; otherwise it returns a failure message. The rationale for not trying to decode when more than t errors have occurred might be ‘we can’t guarantee that we can correct more than t errors, so we won’t bother trying – who would be interested in a decoder that corrects some error patterns of weight greater than t, but not others?’ This defeatist attitude is an example of worst-case-ism, a widespread mental ailment which this book is intended to cure. The fact is that bounded-distance decoders cannot reach the Shannon limit of the binary symmetric channel; only a decoder that often corrects more than t errors can do this. The state of the art in error-correcting codes have decoders that work way beyond the minimum distance of the code.



Definitions of good and bad distance properties Given a family of codes of increasing blocklength N , and with rates approaching a limit R > 0, we may be able to put that family in one of the following categories, which have some similarities to the categories of ‘good’ and ‘bad’ codes defined earlier (p.183): A sequence of codes has ‘good’ distance if d/N tends to a constant greater than zero. A sequence of codes has ‘bad’ distance if d/N tends to zero. A sequence of codes has ‘very bad’ distance if d tends to a constant. Example 13.3. A low-density generator-matrix code is a linear code whose K × N generator matrix G has a small number d 0 of 1s per row, regardless of how big N is. The minimum distance of such a code is at most d 0 , so low-density generator-matrix codes have ‘very bad’ distance. While having large distance is no bad thing, we’ll see, later on, why an emphasis on distance can be unhealthy.

Figure 13.3. The graph of a rate-1/2 low-density generator-matrix code. The rightmost M of the transmitted bits are each connected to a single distinct parity constraint. The leftmost K transmitted bits are each connected to a small number of parity constraints.

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208

13 — Binary Codes Figure 13.4. Schematic picture of part of Hamming space perfectly filled by t-spheres centred on the codewords of a perfect code. t t

... 1

t

2

13.3 Perfect codes A t-sphere (or a sphere of radius t) in Hamming space, centred on a point x, is the set of points whose Hamming distance from x is less than or equal to t. The (7, 4) Hamming code has the beautiful property that if we place 1spheres about each of its 16 codewords, those spheres perfectly fill Hamming space without overlapping. As we saw in Chapter 1, every binary vector of length 7 is within a distance of t = 1 of exactly one codeword of the Hamming code. A code is a perfect t-error-correcting code if the set of t-spheres centred on the codewords of the code fill the Hamming space without overlapping. (See figure 13.4.) Let’s recap our cast of characters. The number of codewords is S = 2 K . The number of points in the entire Hamming space is 2 N . The number of points in a Hamming sphere of radius t is t   X N

w=0

w

.

(13.1)

For a code to be perfect with these parameters, we require S times the number of points in the t-sphere to equal 2N : for a perfect code, 2K or, equivalently,

t   X N

w=0 t  X w=0

w

N w



= 2N

(13.2)

= 2N −K .

(13.3)

For a perfect code, the number of noise vectors in one sphere must equal the number of possible syndromes. The (7, 4) Hamming code satisfies this numerological condition because 1+

  7 = 23 . 1

(13.4)

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13.3: Perfect codes

209

t

... 1

... 1

2

t 2

2

t

... 1

Figure 13.5. Schematic picture of Hamming space not perfectly filled by t-spheres centred on the codewords of a code. The grey regions show points that are at a Hamming distance of more than t from any codeword. This is a misleading picture, as, for any code with large t in high dimensions, the grey space between the spheres takes up almost all of Hamming space.

t

... 1

2

How happy we would be to use perfect codes If there were large numbers of perfect codes to choose from, with a wide range of blocklengths and rates, then these would be the perfect solution to Shannon’s problem. We could communicate over a binary symmetric channel with noise level f , for example, by picking a perfect t-error-correcting code with blocklength N and t = f ∗ N , where f ∗ = f + δ and N and δ are chosen such that the probability that the noise flips more than t bits is satisfactorily small. However, there are almost no perfect codes. The only nontrivial perfect binary codes are 1. the Hamming codes, which are perfect codes with t = 1 and blocklength N = 2M − 1, defined below; the rate of a Hamming code approaches 1 as its blocklength N increases; 2. the repetition codes of odd blocklength N , which are perfect codes with t = (N − 1)/2; the rate of repetition codes goes to zero as 1/N ; and 3. one remarkable 3-error-correcting code with 2 12 codewords of blocklength N = 23 known as the binary Golay code. [A second 2-errorcorrecting Golay code of length N = 11 over a ternary alphabet was discovered by a Finnish football-pool enthusiast called Juhani Virtakallio in 1947.] There are no other binary perfect codes. Why this shortage of perfect codes? Is it because precise numerological coincidences like those satisfied by the parameters of the Hamming code (13.4) and the Golay code, 

     23 23 23 1+ + + = 211 , 1 2 3

(13.5)

are rare? Are there plenty of ‘almost-perfect’ codes for which the t-spheres fill almost the whole space? No. In fact, the picture of Hamming spheres centred on the codewords almost filling Hamming space (figure 13.5) is a misleading one: for most codes, whether they are good codes or bad codes, almost all the Hamming space is taken up by the space between t-spheres (which is shown in grey in figure 13.5). Having established this gloomy picture, we spend a moment filling in the properties of the perfect codes mentioned above.



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210

13 — Binary Codes 00000 0 00000 0 00000 0 00000 0 00000 1 1 1 1 1 1 1 1 1 1 1 1 1 1 1 1 1 1 100000 0 0000 00000 0 1 1 1 1 1 1 1 1 1 1 1 1 1 1 1 1 1 1 1 0000 uN

vN

wN

xN

N

The Hamming codes The (7, 4) Hamming code can be defined as the linear code whose 3 × 7 paritycheck matrix contains, as its columns, all the 7 (= 2 3 − 1) non-zero vectors of length 3. Since these 7 vectors are all different, any single bit-flip produces a distinct syndrome, so all single-bit errors can be detected and corrected. We can generalize this code, with M = 3 parity constraints, as follows. The Hamming codes are single-error-correcting codes defined by picking a number of parity-check constraints, M ; the blocklength N is N = 2 M − 1; the paritycheck matrix contains, as its columns, all the N non-zero vectors of length M bits. The first few Hamming codes have the following rates: Checks, M

(N, K)

R = K/N

2 3 4 5 6

(3, 1) (7, 4) (15, 11) (31, 26) (63, 57)

1/3 4/7 11/15 26/31 57/63

repetition code R3 (7, 4) Hamming code

Exercise 13.4.[2, p.223] What is the probability of block error of the (N, K) Hamming code to leading order, when the code is used for a binary symmetric channel with noise density f ?

13.4 Perfectness is unattainable – first proof We will show in several ways that useful perfect codes do not exist (here, ‘useful’ means ‘having large blocklength N , and rate close neither to 0 nor 1’). Shannon proved that, given a binary symmetric channel with any noise level f , there exist codes with large blocklength N and rate as close as you like to C(f ) = 1 − H2 (f ) that enable communication with arbitrarily small error probability. For large N , the number of errors per block will typically be about fN , so these codes of Shannon are ‘almost-certainly-fN -error-correcting’ codes. Let’s pick the special case of a noisy channel with f ∈ (1/3, 1/2). Can we find a large perfect code that is fN -error-correcting? Well, let’s suppose that such a code has been found, and examine just three of its codewords. (Remember that the code ought to have rate R ' 1 − H 2 (f ), so it should have an enormous number (2N R ) of codewords.) Without loss of generality, we choose one of the codewords to be the all-zero codeword and define the other two to have overlaps with it as shown in figure 13.6. The second codeword differs from the first in a fraction u + v of its coordinates. The third codeword differs from the first in a fraction v + w, and from the second in a fraction u + w. A fraction x of the coordinates have value zero in all three codewords. Now, if the code is fN -error-correcting, its minimum distance must be greater

Figure 13.6. Three codewords.

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13.5: Weight enumerator function of random linear codes

211

than 2fN , so u + v > 2f, v + w > 2f, and u + w > 2f.

(13.6)

Summing these three inequalities and dividing by two, we have u + v + w > 3f.

(13.7)

So if f > 1/3, we can deduce u + v + w > 1, so that x < 0, which is impossible. Such a code cannot exist. So the code cannot have three codewords, let alone 2N R . We conclude that, whereas Shannon proved there are plenty of codes for communicating over a binary symmetric channel with f > 1/3, there are no perfect codes that can do this. We now study a more general argument that indicates that there are no large perfect linear codes for general rates (other than 0 and 1). We do this by finding the typical distance of a random linear code.

N 







































101010100100110100010110 001110111100011001101000 101110111001011000110100 000010111100101101001000 000000110011110100000100 110010001111100000101110 101111100010100001001110 110010110001101011101010 100011100101000010111101 010001000010101001101010 010111110111111110111010 101110101001001101000011

13.5 Weight enumerator function of random linear codes Imagine making a code by picking the binary entries in the M ×N parity-check matrix H at random. What weight enumerator function should we expect? The weight enumerator of one particular code with parity-check matrix H, A(w)H , is the number of codewords of weight w, which can be written A(w)H =

X



[Hx = 0] ,

(13.8)

x:|x|=w

where the sum is over all vectors x whose weight is w and the truth function [Hx = 0] equals one if Hx = 0 and zero otherwise. We can find the expected value of A(w), 

X

hA(w)i =

P (H)A(w)H

(13.9)

H

X X

=

P (H) [Hx = 0] , 

(13.10)

x:|x|=w H

by evaluating the probability that a particular word of weight w > 0 is a codeword of the code (averaging over all binary linear codes in our ensemble). By symmetry, this probability depends only on the weight w of the word, not on the details of the word. The probability that the entire syndrome Hx is zero can be found by multiplying together the probabilities that each of the M bits in the syndrome is zero. Each bit z m of the syndrome is a sum (mod 2) of w random bits, so the probability that z m = 0 is 1/2. The probability that Hx = 0 is thus X (13.11) P (H) [Hx = 0] = (1/2)M = 2−M , 

H

independent of w. The expected number of words of weight w (13.10) is given by summing, over all words of weight w, the probability that each word is a codeword. The  , so number of words of weight w is N w hA(w)i =

  N −M 2 for any w > 0. w

(13.12)

Figure 13.7. A random binary parity-check matrix.





















































M







































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212

13 — Binary Codes

For large N , we can use log

N w



' N H2 (w/N ) and R ' 1 − M/N to write

log2 hA(w)i ' N H2 (w/N ) − M

' N [H2 (w/N ) − (1 − R)] for any w > 0.

(13.13) (13.14)

Gilbert–Varshamov distance For weights w such that H2 (w/N ) < (1 − R), the expectation of A(w) is smaller than 1; for weights such that H 2 (w/N ) > (1 − R), the expectation is greater than 1. We thus expect, for large N , that the minimum distance of a random linear code will be close to the distance d GV defined by (13.15)

Definition. This distance, dGV ≡ N H2−1 (1 − R), is the Gilbert–Varshamov distance for rate R and blocklength N . The Gilbert–Varshamov conjecture, widely believed, asserts that (for large N ) it is not possible to create binary codes with minimum distance significantly greater than dGV . Definition. The Gilbert–Varshamov rate R GV is the maximum rate at which you can reliably communicate with a bounded-distance decoder (as defined on p.207), assuming that the Gilbert–Varshamov conjecture is true.

Why sphere-packing is a bad perspective, and an obsession with distance is inappropriate If one uses a bounded-distance decoder, the maximum tolerable noise level will flip a fraction fbd = 12 dmin /N of the bits. So, assuming dmin is equal to the Gilbert distance dGV (13.15), we have: H2 (2fbd ) = (1 − RGV ).

(13.16)

RGV = 1 − H2 (2fbd ).

(13.17)

Now, here’s the crunch: what did Shannon say is achievable? He said the maximum possible rate of communication is the capacity, C = 1 − H2 (f ).

(13.18)

So for a given rate R, the maximum tolerable noise level, according to Shannon, is given by H2 (f ) = (1 − R). (13.19) Our conclusion: imagine a good code of rate R has been chosen; equations (13.16) and (13.19) respectively define the maximum noise levels tolerable by a bounded-distance decoder, fbd , and by Shannon’s decoder, f . fbd = f /2.

5e+52 4e+52 3e+52 2e+52 1e+52

As a concrete example, figure 13.8 shows the expected weight enumerator function of a rate-1/3 random linear code with N = 540 and M = 360.

H2 (dGV /N ) = (1 − R).

6e+52

(13.20)

Bounded-distance decoders can only ever cope with half the noise-level that Shannon proved is tolerable! How does this relate to perfect codes? A code is perfect if there are tspheres around its codewords that fill Hamming space without overlapping.

0 0

100

0

100

200

300

400

500

1e+60 1e+40 1e+20 1 1e-20 1e-40 1e-60 1e-80 1e-100 1e-120 200

300

400

500

Figure 13.8. The expected weight enumerator function hA(w)i of a random linear code with N = 540 and M = 360. Lower figure shows hA(w)i on a logarithmic scale. 1 Capacity R_GV 0.5

0 0

0.25

0.5

f Figure 13.9. Contrast between Shannon’s channel capacity C and the Gilbert rate RGV – the maximum communication rate achievable using a bounded-distance decoder, as a function of noise level f . For any given rate, R, the maximum tolerable noise level for Shannon is twice as big as the maximum tolerable noise level for a ‘worst-case-ist’ who uses a bounded-distance decoder.

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13.6: Berlekamp’s bats

213

But when a typical random linear code is used to communicate over a binary symmetric channel near to the Shannon limit, the typical number of bits flipped is fN , and the minimum distance between codewords is also fN , or a little bigger, if we are a little below the Shannon limit. So the fN -spheres around the codewords overlap with each other sufficiently that each sphere almost contains the centre of its nearest neighbour! The reason why this overlap is not disastrous is because, in high dimensions, the volume associated with the overlap, shown shaded in figure 13.10, is a tiny fraction of either sphere, so the probability of landing in it is extremely small. The moral of the story is that worst-case-ism can be bad for you, halving your ability to tolerate noise. You have to be able to decode way beyond the minimum distance of a code to get to the Shannon limit! Nevertheless, the minimum distance of a code is of interest in practice, because, under some conditions, the minimum distance dominates the errors made by a code.

Figure 13.10. Two overlapping spheres whose radius is almost as big as the distance between their centres.

13.6 Berlekamp’s bats A blind bat lives in a cave. It flies about the centre of the cave, which corresponds to one codeword, with its typical distance from the centre controlled by a friskiness parameter f . (The displacement of the bat from the centre corresponds to the noise vector.) The boundaries of the cave are made up of stalactites that point in towards the centre of the cave (figure 13.11). Each stalactite is analogous to the boundary between the home codeword and another codeword. The stalactite is like the shaded region in figure 13.10, but reshaped to convey the idea that it is a region of very small volume. Decoding errors correspond to the bat’s intended trajectory passing inside a stalactite. Collisions with stalactites at various distances from the centre are possible. If the friskiness is very small, the bat is usually very close to the centre of the cave; collisions will be rare, and when they do occur, they will usually involve the stalactites whose tips are closest to the centre point. Similarly, under low-noise conditions, decoding errors will be rare, and they will typically involve low-weight codewords. Under low-noise conditions, the minimum distance of a code is relevant to the (very small) probability of error.

t 1

... 2

Figure 13.11. Berlekamp’s schematic picture of Hamming space in the vicinity of a codeword. The jagged solid line encloses all points to which this codeword is the closest. The t-sphere around the codeword takes up a small fraction of this space.

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214

13 — Binary Codes

If the friskiness is higher, the bat may often make excursions beyond the safe distance t where the longest stalactites start, but it will collide most frequently with more distant stalactites, owing to their greater number. There’s only a tiny number of stalactites at the minimum distance, so they are relatively unlikely to cause the errors. Similarly, errors in a real error-correcting code depend on the properties of the weight enumerator function. At very high friskiness, the bat is always a long way from the centre of the cave, and almost all its collisions involve contact with distant stalactites. Under these conditions, the bat’s collision frequency has nothing to do with the distance from the centre to the closest stalactite.

13.7 Concatenation of Hamming codes It is instructive to play some more with the concatenation of Hamming codes, a concept we first visited in figure 11.6, because we will get insights into the notion of good codes and the relevance or otherwise of the minimum distance of a code. We can create a concatenated code for a binary symmetric channel with noise density f by encoding with several Hamming codes in succession. The table recaps the key properties of the Hamming codes, indexed by number of constraints, M . All the Hamming codes have minimum distance d = 3 and can correct one error in N . 2M

N= −1 K =N −M  pB = N3 N2 f 2

blocklength number of source bits probability of block error to leading order

1 0.8

R 0.6 0.4 0.2 0 0

2

4

6

C

8

10

12

Figure 13.12. The rate R of the concatenated Hamming code as a function of the number of concatenations, C.

If we make a product code by concatenating a sequence of C Hamming codes with increasing M , we can choose those parameters {M c }C c=1 in such a way that the rate of the product code

1e+25 1e+20 1e+15

RC =

C Y Nc − M c

c=1

Nc

(13.21)

tends to a non-zero limit as C increases. For example, if we set M 1 = 2, M2 = 3, M3 = 4, etc., then the asymptotic rate is 0.093 (figure 13.12). The blocklength N is a rapidly-growing function of C, so these codes are somewhat impractical. A further weakness of these codes is that their minimum distance is not very good (figure 13.13). Every one of the constituent Hamming codes has minimum distance 3, so the minimum distance of the Cth product is 3C . The blocklength N grows faster than 3 C , so the ratio d/N tends to zero as C increases. In contrast, for typical random codes, the ratio d/N tends to a constant such that H2 (d/N ) = 1 − R. Concatenated Hamming codes thus have ‘bad’ distance. Nevertheless, it turns out that this simple sequence of codes yields good codes for some channels – but not very good codes (see section 11.4 to recall the definitions of the terms ‘good’ and ‘very good’). Rather than prove this result, we will simply explore it numerically. Figure 13.14 shows the bit error probability p b of the concatenated codes assuming that the constituent codes are decoded in sequence, as described in section 11.4. [This one-code-at-a-time decoding is suboptimal, as we saw there.] The horizontal axis shows the rates of the codes. As the number of concatenations increases, the rate drops to 0.093 and the error probability drops towards zero. The channel assumed in the figure is the binary symmetric

1e+10 100000 1 0

2

4

6

C

8

10

12

Figure 13.13. The blocklength NC (upper curve) and minimum distance dC (lower curve) of the concatenated Hamming code as a function of the number of concatenations C.

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13.8: Distance isn’t everything

215

channel with f = 0.0588. This is the highest noise level that can be tolerated using this concatenated code. The take-home message from this story is distance isn’t everything. The minimum distance of a code, although widely worshipped by coding theorists, is not of fundamental importance to Shannon’s mission of achieving reliable communication over noisy channels.

1

pb 0.01

21 N=3 315 61525

0.0001 1e-06 1e-08 1e-10 1e-12

10^13

1e-14 0

0.2

0.4

0.6

R

[3 ]

. Exercise 13.5. Prove that there exist families of codes with ‘bad’ distance that are ‘very good’ codes.

13.8 Distance isn’t everything Let’s get a quantitative feeling for the effect of the minimum distance of a code, for the special case of a binary symmetric channel.

The error probability associated with one low-weight codeword Let a binary code have blocklength N and just two codewords, which differ in d places. For simplicity, let’s assume d is even. What is the error probability if this code is used on a binary symmetric channel with noise level f ? Bit flips matter only in places where the two codewords differ. The error probability is dominated by the probability that d/2 of these bits are flipped. What happens to the other bits is irrelevant, since the optimal decoder ignores them.   d f d/2 (1 − f )d/2 . (13.22) P (block error) ' d/2 This error probability associated with a single codeword of weight d is plotted in figure 13.15. Using the approximation for the binomial coefficient (1.16), we can further approximate P (block error) '

h

2f 1/2 (1 − f )1/2

≡ [β(f )]d ,

id

(13.23) (13.24)

where β(f ) = 2f 1/2 (1 − f )1/2 is called the Bhattacharyya parameter of the channel. Now, consider a general linear code with distance d. Its block error probd  d/2 ability must be at least d/2 f (1 − f )d/2 , independent of the blocklength N of the code. For this reason, a sequence of codes of increasing blocklength N and constant distance d (i.e., ‘very bad’ distance) cannot have a block error probability that tends to zero, on any binary symmetric channel. If we are interested in making superb error-correcting codes with tiny, tiny error probability, we might therefore shun codes with bad distance. However, being pragmatic, we should look more carefully at figure 13.15. In Chapter 1 we argued that codes for disk drives need an error probability smaller than about 10−18 . If the raw error probability in the disk drive is about 0.001, the error probability associated with one codeword at distance d = 20 is smaller than 10−24 . If the raw error probability in the disk drive is about 0.01, the error probability associated with one codeword at distance d = 30 is smaller than 10−20 . For practical purposes, therefore, it is not essential for a code to have good distance. For example, codes of blocklength 10 000, known to have many codewords of weight 32, can nevertheless correct errors of weight 320 with tiny error probability.

0.8

1

Figure 13.14. The bit error probabilities versus the rates R of the concatenated Hamming codes, for the binary symmetric channel with f = 0.0588. Labels alongside the points show the blocklengths, N . The solid line shows the Shannon limit for this channel. The bit error probability drops to zero while the rate tends to 0.093, so the concatenated Hamming codes are a ‘good’ code family.

1 1e-05 1e-10

d=10 d=20 d=30 d=40 d=50 d=60

1e-15 1e-20 0.0001

0.001

0.01

0.1

Figure 13.15. The error probability associated with a single codeword of weight d,  d/2 d (1 − f )d/2 , as a function d/2 f of f .

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216

13 — Binary Codes

I wouldn’t want you to think I am recommending the use of codes with bad distance; in Chapter 47 we will discuss low-density parity-check codes, my favourite codes, which have both excellent performance and good distance.

13.9 The union bound The error probability of a code on the binary symmetric channel can be bounded in terms of its weight enumerator function by adding up appropriate multiples of the error probability associated with a single codeword (13.24): X P (block error) ≤ A(w)[β(f )]w . (13.25) w>0

This inequality, which is an example of a union bound, is accurate for low noise levels f , but inaccurate for high noise levels, because it overcounts the contribution of errors that cause confusion with more than one codeword at a time. . Exercise 13.6.[3 ] Poor man’s noisy-channel coding theorem. Pretending that the union bound (13.25) is accurate, and using the average weight enumerator function of a random linear code (13.14) (section 13.5) as A(w), estimate the maximum rate R UB (f ) at which one can communicate over a binary symmetric channel. Or, to look at it more positively, using the union bound (13.25) as an inequality, show that communication at rates up to R UB (f ) is possible over the binary symmetric channel. In the following chapter, by analysing the probability of error of syndrome decoding for a binary linear code, and using a union bound, we will prove Shannon’s noisy-channel coding theorem (for symmetric binary channels), and thus show that very good linear codes exist.

13.10 Dual codes A concept that has some importance in coding theory, though we will have no immediate use for it in this book, is the idea of the dual of a linear errorcorrecting code. An (N, K) linear error-correcting code can be thought of as a set of 2 K codewords generated by adding together all combinations of K independent basis codewords. The generator matrix of the code consists of those K basis codewords, conventionally written as row vectors. For example, the (7, 4) Hamming code’s generator matrix (from p.10) is   1 0 0 0 1 0 1  0 1 0 0 1 1 0   G= (13.26)  0 0 1 0 1 1 1  0 0 0 1 0 1 1 and its sixteen codewords were displayed in table 1.14 (p.9). The codewords of this code are linear combinations of the four vectors [1 0 0 0 1 0 1], [0 1 0 0 1 1 0], [0 0 1 0 1 1 1], and [0 0 0 1 0 1 1]. An (N, K) code may also be described in terms of an M × N parity-check matrix (where M = N − K) as the set of vectors {t} that satisfy Ht = 0.

(13.27)

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13.10: Dual codes

217

One way of thinking of this equation is that each row of H specifies a vector to which t must be orthogonal if it is a codeword. The generator matrix specifies K vectors from which all codewords can be built, and the parity-check matrix specifies a set of M vectors to which all codewords are orthogonal. The dual of a code is obtained by exchanging the generator matrix and the parity-check matrix. Definition. The set of all vectors of length N that are orthogonal to all codewords in a code, C, is called the dual of the code, C ⊥ . If t is orthogonal to h1 and h2 , then it is also orthogonal to h3 ≡ h1 + h2 ; so all codewords are orthogonal to any linear combination of the M rows of H. So the set of all linear combinations of the rows of the parity-check matrix is the dual code. For our Hamming (7, 4) code, the parity-check matrix is (from p.12):

H=



P I3





 1 1 1 0 1 0 0 =  0 1 1 1 0 1 0 . 1 0 1 1 0 0 1

(13.28)

The dual of the (7, 4) Hamming code H(7,4) is the code shown in table 13.16. 0000000 0010111

0101101 0111010

1001110 1011001

Table 13.16. The eight codewords of the dual of the (7, 4) Hamming code. [Compare with table 1.14, p.9.]

1100011 1110100

A possibly unexpected property of this pair of codes is that the dual, ⊥ , is contained within the code H H(7,4) (7,4) itself: every word in the dual code is a codeword of the original (7, 4) Hamming code. This relationship can be written using set notation: ⊥ H(7,4) ⊂ H(7,4) .

(13.29)

The possibility that the set of dual vectors can overlap the set of codeword vectors is counterintuitive if we think of the vectors as real vectors – how can a vector be orthogonal to itself? But when we work in modulo-two arithmetic, many non-zero vectors are indeed orthogonal to themselves! . Exercise 13.7.[1, p.223] Give a simple rule that distinguishes whether a binary vector is orthogonal to itself, as is each of the three vectors [1 1 1 0 1 0 0], [0 1 1 1 0 1 0], and [1 0 1 1 0 0 1].

Some more duals In general, if a code has a systematic generator matrix, G = [IK |PT] ,

(13.30)

where P is a K × M matrix, then its parity-check matrix is H = [P|IM ] .

(13.31)

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218

13 — Binary Codes

Example 13.8. The repetition code R 3 has generator matrix   G= 1 1 1 ;

(13.32)

its parity-check matrix is

H=



1 1 0 1 0 1



.

(13.33)

The two codewords are [1 1 1] and [0 0 0]. The dual code has generator matrix   1 1 0 G⊥ = H = 1 0 1

(13.34)

or equivalently, modifying G⊥ into systematic form by row additions,   1 0 1 ⊥ G = . (13.35) 0 1 1 We call this dual code the simple parity code P 3 ; it is the code with one parity-check bit, which is equal to the sum of the two source bits. The dual code’s four codewords are [1 1 0], [1 0 1], [0 0 0], and [0 1 1]. In this case, the only vector common to the code and the dual is the all-zero codeword.

Goodness of duals If a sequence of codes is ‘good’, are their duals good too? Examples can be constructed of all cases: good codes with good duals (random linear codes); bad codes with bad duals; and good codes with bad duals. The last category is especially important: many state-of-the-art codes have the property that their duals are bad. The classic example is the low-density parity-check code, whose dual is a low-density generator-matrix code. . Exercise 13.9.[3 ] Show that low-density generator-matrix codes are bad. A family of low-density generator-matrix codes is defined by two parameters j, k, which are the column weight and row weight of all rows and columns respectively of G. These weights are fixed, independent of N ; for example, (j, k) = (3, 6). [Hint: show that the code has low-weight codewords, then use the argument from p.215.] Exercise 13.10. [5 ] Show that low-density parity-check codes are good, and have good distance. (For solutions, see Gallager (1963) and MacKay (1999b).)

Self-dual codes The (7, 4) Hamming code had the property that the dual was contained in the code itself. A code is self-orthogonal if it is contained in its dual. For example, the dual of the (7, 4) Hamming code is a self-orthogonal code. One way of seeing this is that the overlap between any pair of rows of H is even. Codes that contain their duals are important in quantum error-correction (Calderbank and Shor, 1996). It is intriguing, though not necessarily useful, to look at codes that are self-dual. A code C is self-dual if the dual of the code is identical to the code. C ⊥ = C.

Some properties of self-dual codes can be deduced:

(13.36)

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13.11: Generalizing perfectness to other channels

219

1. If a code is self-dual, then its generator matrix is also a parity-check matrix for the code. 2. Self-dual codes have rate 1/2, i.e., M = K = N/2. 3. All codewords have even weight. . Exercise 13.11. [2, p.223] What property must the matrix P satisfy, if the code with generator matrix G = [IK |PT] is self-dual? Examples of self-dual codes 1. The repetition code R2 is a simple example of a self-dual code.   G=H= 1 1 . (13.37) 2. The smallest non-trivial self-dual code is  1 0 0 0    0 1 0 0 G = I4 PT =   0 0 1 0 0 0 0 1

the following (8, 4) code.  0 1 1 1 1 0 1 1  . (13.38) 1 1 0 1  1 1 1 0

. Exercise 13.12. [2, p.223] Find the relationship of the above (8, 4) code to the (7, 4) Hamming code.

Duals and graphs Let a code be represented by a graph in which there are nodes of two types, parity-check constraints and equality constraints, joined by edges which represent the bits of the code (not all of which need be transmitted). The dual code’s graph is obtained by replacing all parity-check nodes by equality nodes and vice versa. This type of graph is called a normal graph by Forney (2001).

Further reading Duals are important in coding theory because functions involving a code (such as the posterior distribution over codewords) can be transformed by a Fourier transform into functions over the dual code. For an accessible introduction to Fourier analysis on finite groups, see Terras (1999). See also MacWilliams and Sloane (1977).

13.11 Generalizing perfectness to other channels Having given up on the search for perfect codes for the binary symmetric channel, we could console ourselves by changing channel. We could call a code ‘a perfect u-error-correcting code for the binary erasure channel’ if it can restore any u erased bits, and never more than u. Rather than using the word perfect, however, the conventional term for such a code is a ‘maximum distance separable code’, or MDS code. As we already noted in exercise 11.10 (p.190), the (7, 4) Hamming code is not an MDS code. It can recover some sets of 3 erased bits, but not all. If any 3 bits corresponding to a codeword of weight 3 are erased, then one bit of information is unrecoverable. This is why the (7, 4) code is a poor choice for a RAID system.

In a perfect u-error-correcting code for the binary erasure channel, the number of redundant bits must be N − K = u.

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220

13 — Binary Codes

A tiny example of a maximum distance separable code is the simple paritycheck code P3 whose parity-check matrix is H = [1 1 1]. This code has 4 codewords, all of which have even parity. All codewords are separated by a distance of 2. Any single erased bit can be restored by setting it to the parity of the other two bits. The repetition codes are also maximum distance separable codes. . Exercise 13.13. [5, p.224] Can you make an (N, K) code, with M = N − K parity symbols, for a q-ary erasure channel, such that the decoder can recover the codeword when any M symbols are erased in a block of N ? [Example: for the channel with q = 4 symbols there is an (N, K) = (5, 2) code which can correct any M = 3 erasures.] For the q-ary erasure channel with q > 2, there are large numbers of MDS codes, of which the Reed–Solomon codes are the most famous and most widely used. As long as the field size q is bigger than the blocklength N , MDS block codes of any rate can be found. (For further reading, see Lin and Costello (1983).)

13.12 Summary Shannon’s codes for the binary symmetric channel can almost always correct fN errors, but they are not fN -error-correcting codes.

Reasons why the distance of a code has little relevance 1. The Shannon limit shows that the best codes must be able to cope with a noise level twice as big as the maximum noise level for a boundeddistance decoder. 2. When the binary symmetric channel has f > 1/4, no code with a bounded-distance decoder can communicate at all; but Shannon says good codes exist for such channels. 3. Concatenation shows that we can get good performance even if the distance is bad. The whole weight enumerator function is relevant to the question of whether a code is a good code. The relationship between good codes and distance properties is discussed further in exercise 13.14 (p.220).

13.13 Further exercises Exercise 13.14. [3, p.224] A codeword t is selected from a linear (N, K) code C, and it is transmitted over a noisy channel; the received signal is y. We assume that the channel is a memoryless channel such as a Gaussian channel. Given an assumed channel model P (y | t), there are two decoding problems. The codeword decoding problem is the task of inferring which codeword t was transmitted given the received signal. The bitwise decoding problem is the task of inferring for each transmitted bit tn how likely it is that that bit was a one rather than a zero.

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13.13: Further exercises

221

Consider optimal decoders for these two decoding problems. Prove that the probability of error of the optimal bitwise-decoder is closely related to the probability of error of the optimal codeword-decoder, by proving the following theorem. Theorem 13.1 If a binary linear code has minimum distance d min , then, for any given channel, the codeword bit error probability of the optimal bitwise decoder, pb , and the block error probability of the maximum likelihood decoder, pB , are related by: pB ≥ p b ≥

1 dmin pB . 2 N

(13.39)

Exercise 13.15. [1 ] What are the minimum distances of the (15, 11) Hamming code and the (31, 26) Hamming code? . Exercise 13.16. [2 ] Let A(w) be the average weight enumerator function of a rate-1/3 random linear code with N = 540 and M = 360. Estimate, from first principles, the value of A(w) at w = 1. Exercise 13.17. [3C ] A code with minimum distance greater than d GV . A rather nice (15, 5) code is generated by this generator matrix, which is based on measuring the parities of all the 53 = 10 triplets of source bits:   1 · · · · · 1 1 1 · · 1 1 · 1  · 1 · · · · · 1 1 1 1 · 1 1 ·     (13.40) G=  · · 1 · · 1 · · 1 1 · 1 · 1 1 .  · · · 1 · 1 1 · · 1 1 · 1 · 1  · · · · 1 1 1 1 · · 1 1 · 1 · Find the minimum distance and weight enumerator function of this code.

Exercise 13.18. [3C ] Find the minimum distance of the ‘pentagonful’ lowdensity parity-check code whose parity-check matrix is   1 · · · 1 1 · · · · · · · · ·  1 1 · · · · 1 · · · · · · · ·     · 1 1 · · · · 1 · · · · · · ·     · · 1 1 · · · · 1 · · · · · ·     · · · 1 1 · · · · 1 · · · · ·   . H= (13.41)   · · · · · 1 · · · · 1 · · · 1   · · · · · · · · 1 · 1 1 · · ·     · · · · · · 1 · · · · 1 1 · ·     · · · · · · · · · 1 · · 1 1 ·  · · · · · · · 1 · · · · · 1 1 Show that nine of the ten rows are independent, so the code has parameters N = 15, K = 6. Using a computer, find its weight enumerator function.

. Exercise 13.19. [3C ] Replicate the calculations used to produce figure 13.12. Check the assertion that the highest noise level that’s correctable is 0.0588. Explore alternative concatenated sequences of codes. Can you find a better sequence of concatenated codes – better in the sense that it has either higher asymptotic rate R or can tolerate a higher noise level f?

Figure 13.17. The graph of the pentagonful low-density parity-check code with 15 bit nodes (circles) and 10 parity-check nodes (triangles). [This graph is known as the Petersen graph.]

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222

13 — Binary Codes

Exercise 13.20. [3, p.226] Investigate the possibility of achieving the Shannon limit with linear block codes, using the following counting argument. Assume a linear code of large blocklength N and rate R = K/N . The code’s parity-check matrix H has M = N − K rows. Assume that the code’s optimal decoder, which solves the syndrome decoding problem Hn = z, allows reliable communication over a binary symmetric channel with flip probability f . How many ‘typical’ noise vectors n are there? Roughly how many distinct syndromes z are there? Since n is reliably deduced from z by the optimal decoder, the number of syndromes must be greater than or equal to the number of typical noise vectors. What does this tell you about the largest possible value of rate R for a given f ? . Exercise 13.21. [2 ] Linear binary codes use the input symbols 0 and 1 with equal probability, implicitly treating the channel as a symmetric channel. Investigate how much loss in communication rate is caused by this assumption, if in fact the channel is a highly asymmetric channel. Take as an example a Z-channel. How much smaller is the maximum possible rate of communication using symmetric inputs than the capacity of the channel? [Answer: about 6%.] Exercise 13.22. [2 ] Show that codes with ‘very bad’ distance are ‘bad’ codes, as defined in section 11.4 (p.183). Exercise 13.23. [3 ] One linear code can be obtained from another by puncturing. Puncturing means taking each codeword and deleting a defined set of bits. Puncturing turns an (N, K) code into an (N 0 , K) code, where N0 < N. Another way to make new linear codes from old is shortening. Shortening means constraining a defined set of bits to be zero, and then deleting them from the codewords. Typically if we shorten by one bit, half of the code’s codewords are lost. Shortening typically turns an (N, K) code into an (N 0 , K 0 ) code, where N − N 0 = K − K 0 . Another way to make a new linear code from two old ones is to make the intersection of the two codes: a codeword is only retained in the new code if it is present in both of the two old codes. Discuss the effect on a code’s distance-properties of puncturing, shortening, and intersection. Is it possible to turn a code family with bad distance into a code family with good distance, or vice versa, by each of these three manipulations? Exercise 13.24. [3, p.226] Todd Ebert’s ‘hat puzzle’. Three players enter a room and a red or blue hat is placed on each person’s head. The colour of each hat is determined by a coin toss, with the outcome of one coin toss having no effect on the others. Each person can see the other players’ hats but not his own. No communication of any sort is allowed, except for an initial strategy session before the group enters the room. Once they have had a chance to look at the other hats, the players must simultaneously guess their

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13.14: Solutions

223

own hat’s colour or pass. The group shares a $3 million prize if at least one player guesses correctly and no players guess incorrectly. The same game can be played with any number of players. The general problem is to find a strategy for the group that maximizes its chances of winning the prize. Find the best strategies for groups of size three and seven. [Hint: when you’ve done three and seven, you might be able to solve fifteen.] Exercise 13.25. [5 ] Estimate how many binary low-density parity-check codes have self-orthogonal duals. [Note that we don’t expect a huge number, since almost all low-density parity-check codes are ‘good’, but a lowdensity parity-check code that contains its dual must be ‘bad’.] Exercise 13.26. [2C ] In figure 13.15 we plotted the error probability associated with a single codeword of weight d as a function of the noise level f of a binary symmetric channel. Make an equivalent plot for the case of the Gaussian channel, showing the error probability associated with a single codeword of weight d as a function of the rate-compensated signal-tonoise ratio Eb /N0 . Because Eb /N0 depends on the rate, you have to choose a code rate. Choose R = 1/2, 2/3, 3/4, or 5/6.

13.14 Solutions Solution to exercise13.4 (p.210). order is pB = N3 N2 f 2 .

The probability of block error to leading

Solution to exercise 13.7 (p.217). A binary vector is perpendicular to itself if it has even weight, i.e., an even number of 1s. Solution to exercise 13.11 (p.219). The self-dual code has two equivalent parity-check matrices, H1 = G = [IK |PT] and H2 = [P|IK ]; these must be equivalent to each other through row additions, that is, there is a matrix U such that UH2 = H1 , so [UP|UIK ] = [IK |PT] .

(13.42)

From the right-hand sides of this equation, we have U = PT, so the left-hand sides become: PTP = IK . (13.43) Thus if a code with generator matrix G = [I K |PT] is self-dual then P is an orthogonal matrix, modulo 2, and vice versa. Solution to exercise 13.12 (p.219). The (8, 4) and (7, 4) related. The (8, 4) code, whose parity-check matrix is  0 1 1 1 1 0 0   1 0 1 1 0 1 0  H = P I4 =   1 1 0 1 0 0 1 1 1 1 0 0 0 0

codes are intimately  0 0  , 0  1

(13.44)

is obtained by (a) appending an extra parity-check bit which can be thought of as the parity of all seven bits of the (7, 4) Hamming code; and (b) reordering the first four bits.

If you already know the hat puzzle, you could try the ‘Scottish version’ of the rules in which the prize is only awarded to the group if they all guess correctly. In the ‘Reformed Scottish version’, all the players must guess correctly, and there are two rounds of guessing. Those players who guess during round one leave the room. The remaining players must guess in round two. What strategy should the team adopt to maximize their chance of winning?

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224

13 — Binary Codes

Solution to exercise 13.13 (p.220). If an (N, K) code, with M = N − K parity symbols, has the property that the decoder can recover the codeword when any M symbols are erased in a block of N , then the code is said to be maximum distance separable (MDS). No MDS binary codes exist, apart from the repetition codes and simple parity codes. For q > 2, some MDS codes can be found. As a simple example, here is a (9, 2) code for the 8-ary erasure channel. The code is defined in terms of the multiplication and addition rules of GF (8), which are given in Appendix C.1. The elements of the input alphabet are {0, 1, A, B, C, D, E, F } and the generator matrix of the code is   1 0 1 A B C D E F . (13.45) G= 0 1 1 1 1 1 1 1 1 The resulting 64 codewords are: 000000000 101ABCDEF A0ACEB1FD B0BEDFC1A C0CBFEAD1 D0D1CAFBE E0EF1DBAC F0FDA1ECB

011111111 110BADCFE A1BDFA0EC B1AFCED0B C1DAEFBC0 D1C0DBEAF E1FE0CABD F1ECB0FDA

0AAAAAAAA 1AB01EFCD AA0EC1BDF BA1CFDEB0 CAE1DC0FB DAFBE0D1C EACDBF10E FADF0BCE1

0BBBBBBBB 1BA10FEDC AB1FD0ACE BB0DECFA1 CBF0CD1EA DBEAF1C0D EBDCAE01F FBCE1ADF0

0CCCCCCCC 1CDEF01AB ACE0AFDB1 BCFA1B0DE CC0FBAE1D DC1D0EBFA ECABD1FE0 FCB1EDA0F

0DDDDDDDD 1DCFE10BA ADF1BECA0 BDEB0A1CF CD1EABF0C DD0C1FAEB EDBAC0EF1 FDA0FCB1E

Solution to exercise 13.14 (p.220). Quick, rough proof of the theorem. Let x denote the difference between the reconstructed codeword and the transmitted codeword. For any given channel output r, there is a posterior distribution over x. This posterior distribution is positive only on vectors x belonging to the code; the sums that follow are over codewords x. The block error probability is: X pB = P (x | r). (13.46) x6=0

The average bit error probability, averaging over all bits in the codeword, is: pb =

X

x6=0

P (x | r)

w(x) , N

(13.47)

where w(x) is the weight of codeword x. Now the weights of the non-zero codewords satisfy w(x) dmin ≥ . (13.48) 1≥ N N Substituting the inequalities (13.48) into the definitions (13.46, 13.47), we obtain: dmin pB ≥ p b ≥ pB , (13.49) N which is a factor of two stronger, on the right, than the stated result (13.39). In making the proof watertight, I have weakened the result a little. Careful proof. The theorem relates the performance of the optimal block decoding algorithm and the optimal bitwise decoding algorithm. We introduce another pair of decoding algorithms, called the blockguessing decoder and the bit-guessing decoder. The idea is that these two algorithms are similar to the optimal block decoder and the optimal bitwise decoder, but lend themselves more easily to analysis. We now define these decoders. Let x denote the inferred codeword. For any given code:

0EEEEEEEE 1EFCDAB01 AECA0DF1B BED0B1AFC CEAD10CBF DEBFAC1D0 EE01FBDCA FE1BCF0AD

0FFFFFFFF 1FEDCBA10 AFDB1CE0A BFC1A0BED CFBC01DAE DFAEBD0C1 EF10EACDB FF0ADE1BC

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13.14: Solutions

225

The optimal block decoder returns the codeword x that maximizes the posterior probability P (x | r), which is proportional to the likelihood P (r | x). The probability of error of this decoder is called p B .

The optimal bit decoder returns for each of the N bits, x n , the value of a that maximizes the posterior probability P (x n = a | r) = P x P (x | r) [xn = a]. 

The probability of error of this decoder is called p b .

The block-guessing decoder returns a random codeword x with probability distribution given by the posterior probability P (x | r). The probability of error of this decoder is called p G B.

The bit-guessing decoder returns for each of the N bits, x n , a random bit from the probability distribution P (x n = a | r). The probability of error of this decoder is called p G b.

The theorem states that the optimal bit error probability p b is bounded above by pB and below by a given multiple of pB (13.39). The left-hand inequality in (13.39) is trivially true – if a block is correct, all its constituent bits are correct; so if the optimal block decoder outperformed the optimal bit decoder, we could make a better bit decoder from the block decoder. We prove the right-hand inequality by establishing that: (a) the bit-guessing decoder is nearly as good as the optimal bit decoder: pG b ≤ 2pb .

(13.50)

(b) the bit-guessing decoder’s error probability is related to the blockguessing decoder’s by dmin G pG p . (13.51) b ≥ N B Then since pG B ≥ pB , we have 1 dmin G 1 dmin 1 p ≥ pB . ≥ pb > pG 2 b 2 N B 2 N

(13.52)

We now prove the two lemmas. Near-optimality of guessing: Consider first the case of a single bit, with posterior probability {p0 , p1 }. The optimal bit decoder has probability of error P optimal = min(p0 , p1 ).

(13.53)

The guessing decoder picks from 0 and 1. The truth is also distributed with the same probability. The probability that the guesser and the truth match is p20 + p21 ; the probability that they mismatch is the guessing error probability, P guess = 2p0 p1 ≤ 2 min(p0 , p1 ) = 2P optimal .

(13.54)

guess , and p is the Since pG b b is the average of many such error probabilities, P average of the corresponding optimal error probabilities, P optimal , we obtain the desired relationship (13.50) between p G 2 b and pb .

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226

13 — Binary Codes

Relationship between bit error probability and block error probability: The bitguessing and block-guessing decoders can be combined in a single system: we can draw a sample xn from the marginal distribution P (x n | r) by drawing a sample (xn , x) from the joint distribution P (x n , x | r), then discarding the value of x. We can distinguish between two cases: the discarded value of x is the correct codeword, or not. The probability of bit error for the bit-guessing decoder can then be written as a sum of two terms: pG b

= =

P (x correct)P (bit error | x correct)

+ P (x incorrect)P (bit error | x incorrect)

0 + pG B P (bit error | x incorrect).

Now, whenever the guessed x is incorrect, the true x must differ from it in at least d bits, so the probability of bit error in these cases is at least d/N . So pG b ≥

d G p . N B

QED.

2

Solution to exercise 13.20 (p.222). The number of ‘typical’ noise vectors n is roughly 2N H2 (f ) . The number of distinct syndromes z is 2 M . So reliable communication implies M ≥ N H2 (f ), (13.55) or, in terms of the rate R = 1 − M/N , R ≤ 1 − H2 (f ),

(13.56)

a bound which agrees precisely with the capacity of the channel. This argument is turned into a proof in the following chapter. Solution to exercise 13.24 (p.222). In the three-player case, it is possible for the group to win three-quarters of the time. Three-quarters of the time, two of the players will have hats of the same colour and the third player’s hat will be the opposite colour. The group can win every time this happens by using the following strategy. Each player looks at the other two players’ hats. If the two hats are different colours, he passes. If they are the same colour, the player guesses his own hat is the opposite colour. This way, every time the hat colours are distributed two and one, one player will guess correctly and the others will pass, and the group will win the game. When all the hats are the same colour, however, all three players will guess incorrectly and the group will lose. When any particular player guesses a colour, it is true that there is only a 50:50 chance that their guess is right. The reason that the group wins 75% of the time is that their strategy ensures that when players are guessing wrong, a great many are guessing wrong. For larger numbers of players, the aim is to ensure that most of the time no one is wrong and occasionally everyone is wrong at once. In the game with 7 players, there is a strategy for which the group wins 7 out of every 8 times they play. In the game with 15 players, the group can win 15 out of 16 times. If you have not figured out these winning strategies for teams of 7 and 15, I recommend thinking about the solution to the three-player game in terms

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13.14: Solutions of the locations of the winning and losing states on the three-dimensional hypercube, then thinking laterally. If the number of players, N , is 2r − 1, the optimal strategy can be defined using a Hamming code of length N , and the probability of winning the prize is N/(N + 1). Each player is identified with a number n ∈ 1 . . . N . The two colours are mapped onto 0 and 1. Any state of their hats can be viewed as a received vector out of a binary channel. A random binary vector of length N is either a codeword of the Hamming code, with probability 1/(N + 1), or it differs in exactly one bit from a codeword. Each player looks at all the other bits and considers whether his bit can be set to a colour such that the state is a codeword (which can be deduced using the decoder of the Hamming code). If it can, then the player guesses that his hat is the other colour. If the state is actually a codeword, all players will guess and will guess wrong. If the state is a non-codeword, only one player will guess, and his guess will be correct. It’s quite easy to train seven players to follow the optimal strategy if the cyclic representation of the (7, 4) Hamming code is used (p.19).

227

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About Chapter 14 In this chapter we will draw together several ideas that we’ve encountered so far in one nice short proof. We will simultaneously prove both Shannon’s noisy-channel coding theorem (for symmetric binary channels) and his source coding theorem (for binary sources). While this proof has connections to many preceding chapters in the book, it’s not essential to have read them all. On the noisy-channel coding side, our proof will be more constructive than the proof given in Chapter 10; there, we proved that almost any random code is ‘very good’. Here we will show that almost any linear code is very good. We will make use of the idea of typical sets (Chapters 4 and 10), and we’ll borrow from the previous chapter’s calculation of the weight enumerator function of random linear codes (section 13.5). On the source coding side, our proof will show that random linear hash functions can be used for compression of compressible binary sources, thus giving a link to Chapter 12.

228

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14 Very Good Linear Codes Exist In this chapter we’ll use a single calculation to prove simultaneously the source coding theorem and the noisy-channel coding theorem for the binary symmetric channel. Incidentally, this proof works for much more general channel models, not only the binary symmetric channel. For example, the proof can be reworked for channels with non-binary outputs, for time-varying channels and for channels with memory, as long as they have binary inputs satisfying a symmetry property, cf. section 10.6.

14.1 A simultaneous proof of the source coding and noisy-channel coding theorems We consider a linear error-correcting code with binary parity-check matrix H. The matrix has M rows and N columns. Later in the proof we will increase N and M , keeping M ∝ N . The rate of the code satisfies R≥1−

M . N

(14.1)

If all the rows of H are independent then this is an equality, R = 1 − M/N . In what follows, we’ll assume the equality holds. Eager readers may work out the expected rank of a random binary matrix H (it’s very close to M ) and pursue the effect that the difference (M − rank) has on the rest of this proof (it’s negligible). A codeword t is selected, satisfying Ht = 0 mod 2,

(14.2)

and a binary symmetric channel adds noise x, giving the received signal r = t + x mod 2.

(14.3)

The receiver aims to infer both t and x from r using a syndrome-decoding approach. Syndrome decoding was first introduced in section 1.2 (p.10 and 11). The receiver computes the syndrome z = Hr mod 2 = Ht + Hx mod 2 = Hx mod 2.

(14.4)

The syndrome only depends on the noise x, and the decoding problem is to find the most probable x that satisfies Hx = z mod 2. 229

(14.5)

In this chapter x denotes the noise added by the channel, not the input to the channel.

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230

14 — Very Good Linear Codes Exist

ˆ , is then subtracted from r to give the This best estimate for the noise vector, x best guess for t. Our aim is to show that, as long as R < 1−H(X) = 1−H 2 (f ), where f is the flip probability of the binary symmetric channel, the optimal decoder for this syndrome-decoding problem has vanishing probability of error, as N increases, for random H. We prove this result by studying a sub-optimal strategy for solving the decoding problem. Neither the optimal decoder nor this typical-set decoder would be easy to implement, but the typical-set decoder is easier to analyze. The typical-set decoder examines the typical set T of noise vectors, the set of noise vectors x0 that satisfy log 1/P (x0 ) ' N H(X), checking to see if any of those typical vectors x0 satisfies the observed syndrome, Hx0 = z.

(14.6)

If exactly one typical vector x0 does so, the typical set decoder reports that vector as the hypothesized noise vector. If no typical vector matches the observed syndrome, or more than one does, then the typical set decoder reports an error. The probability of error of the typical-set decoder, for a given matrix H, can be written as a sum of two terms, (II)

PTS|H = P (I) + PTS|H ,

We’ll leave out the s and βs that make a typical-set definition rigorous. Enthusiasts are encouraged to revisit section 4.4 and put these details into this proof.

(14.7)

where P (I) is the probability that the true noise vector x is itself not typical, (II) and PTS|H is the probability that the true x is typical and at least one other typical vector clashes with it. The first probability vanishes as N increases, as we proved when we first studied typical sets (Chapter 4). We concentrate on the second probability. To recap, we’re imagining a true noise vector, x; and if any of the typical noise vectors x 0 , different from x, satisfies H(x0 − x) = 0, then we have an error. We use the truth function   H(x0 − x) = 0 , (14.8) 

whose value is one if the statement H(x 0 − x) = 0 is true and zero otherwise. We can bound the number of type II errors made when the noise is x thus: X   H(x0 − x) = 0 . (14.9) [Number of errors given x and H] ≤ 

x0 :

x0 ∈ T x0 6= x

The number of errors is either zero or one; the sum on the right-hand side may exceed one, in cases where several typical noise vectors have the same syndrome. We can now write down the probability of a type-II error by averaging over x: X  X  (II) P (x) PTS|H ≤ H(x0 − x) = 0 . (14.10) 

x∈T

x0 :

x0 ∈ T x0 6= x

Now, we will find the average of this probability of type-II error over all linear codes by averaging over H. By showing that the average probability of type-II error vanishes, we will thus show that there exist linear codes with vanishing error probability, indeed, that almost all linear codes are very good. We denote averaging over all binary matrices H by h. . .i H . The average probability of type-II error is E D X (II) (II) (II) P¯TS = (14.11) P (H)PTS|H = PTS|H H

H

Equation (14.9) is a union bound.

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14.2: Data compression by linear hash codes

= =

*

X

P (x)

x∈T

X

x∈T

X 0

x:

P (x)





x0 ∈ T x0 6= x

X  

x0 :

x0 ∈ T x0 6= x

231 0

H(x − x) = 0 0

H(x − x) = 0



+



(14.12)

H

H

.

(14.13)

Now, the quantity h [H(x0 − x) = 0]iH already cropped up when we were calculating the expected weight enumerator function of random linear codes (section 13.5): for any non-zero binary vector v, the probability that Hv = 0, averaging over all matrices H, is 2−M . So ! X (II) ¯ P = P (x) (|T | − 1) 2−M (14.14) 

TS

x∈T

≤ |T | 2−M ,

(14.15)

where |T | denotes the size of the typical set. As you will recall from Chapter 4, there are roughly 2N H(X) noise vectors in the typical set. So (II) P¯TS ≤ 2N H(X) 2−M .

(14.16)

This bound on the probability of error either vanishes or grows exponentially as N increases (remembering that we are keeping M proportional to N as N increases). It vanishes if H(X) < M/N. (14.17) Substituting R = 1 − M/N , we have thus established the noisy-channel coding theorem for the binary symmetric channel: very good linear codes exist for any rate R satisfying R < 1 − H(X), (14.18) where H(X) is the entropy of the channel noise, per bit.

2

Exercise 14.1.[3 ] Redo the proof for a more general channel.

14.2 Data compression by linear hash codes The decoding game we have just played can also be viewed as an uncompression game. The world produces a binary noise vector x from a source P (x). The noise has redundancy (if the flip probability is not 0.5). We compress it with a linear compressor that maps the N -bit input x (the noise) to the M -bit output z (the syndrome). Our uncompression task is to recover the input x from the output z. The rate of the compressor is Rcompressor ≡ M/N.

(14.19)

[We don’t care about the possibility of linear redundancies in our definition of the rate, here.] The result that we just found, that the decoding problem can be solved, for almost any H, with vanishing error probability, as long as H(X) < M/N , thus instantly proves a source coding theorem: Given a binary source X of entropy H(X), and a required compressed rate R > H(X), there exists a linear compressor x → z = Hx mod 2 having rate M/N equal to that required rate R, and an associated uncompressor, that is virtually lossless. This theorem is true not only for a source of independent identically distributed symbols but also for any source for which a typical set can be defined: sources with memory, and time-varying sources, for example; all that’s required is that the source be ergodic.

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232

14 — Very Good Linear Codes Exist

Notes This method for proving that codes are good can be applied to other linear codes, such as low-density parity-check codes (MacKay, 1999b; Aji et al., 2000). For each code we need an approximation of its expected weight enumerator function.

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15 Further Exercises on Information Theory The most exciting exercises, which will introduce you to further ideas in information theory, are towards the end of this chapter.

Refresher exercises on source coding and noisy channels . Exercise 15.1.[2 ] Let X be an ensemble with AX = {0, 1} and PX = {0.995, 0.005}. Consider source coding using the block coding of X 100 where every x ∈ X 100 containing 3 or fewer 1s is assigned a distinct codeword, while the other xs are ignored. (a) If the assigned codewords are all of the same length, find the minimum length required to provide the above set with distinct codewords. (b) Calculate the probability of getting an x that will be ignored. . Exercise 15.2.[2 ] Let X be an ensemble with PX = {0.1, 0.2, 0.3, 0.4}. The ensemble is encoded using the symbol code C = {0001, 001, 01, 1}. Consider the codeword corresponding to x ∈ X N , where N is large. (a) Compute the entropy of the fourth bit of transmission. (b) Compute the conditional entropy of the fourth bit given the third bit. (c) Estimate the entropy of the hundredth bit. (d) Estimate the conditional entropy of the hundredth bit given the ninety-ninth bit. Exercise 15.3.[2 ] Two fair dice are rolled by Alice and the sum is recorded. Bob’s task is to ask a sequence of questions with yes/no answers to find out this number. Devise in detail a strategy that achieves the minimum possible average number of questions. . Exercise 15.4.[2 ] How can you use a coin to draw straws among 3 people? . Exercise 15.5.[2 ] In a magic trick, there are three participants: the magician, an assistant, and a volunteer. The assistant, who claims to have paranormal abilities, is in a soundproof room. The magician gives the volunteer six blank cards, five white and one blue. The volunteer writes a different integer from 1 to 100 on each card, as the magician is watching. The volunteer keeps the blue card. The magician arranges the five white cards in some order and passes them to the assistant. The assistant then announces the number on the blue card. How does the trick work? 233

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234

15 — Further Exercises on Information Theory

. Exercise 15.6.[3 ] How does this trick work? ‘Here’s an ordinary pack of cards, shuffled into random order. Please choose five cards from the pack, any that you wish. Don’t let me see their faces. No, don’t give them to me: pass them to my assistant Esmerelda. She can look at them. ‘Now, Esmerelda, show me four of the cards. Hmm. . . nine of spades, six of clubs, four of hearts, ten of diamonds. The hidden card, then, must be the queen of spades!’ The trick can be performed as described above for a pack of 52 cards. Use information theory to give an upper bound on the number of cards for which the trick can be performed. . Exercise 15.7.[2 ] Find a probability sequence p = (p1 , p2 , . . .) such that H(p) = ∞. . Exercise 15.8.[2 ] Consider a discrete memoryless source with A X = {a, b, c, d} and PX = {1/2, 1/4, 1/8, 1/8}. There are 48 = 65 536 eight-letter words that can be formed from the four letters. Find the total number of such words that are in the typical set TN β (equation 4.29) where N = 8 and β = 0.1. . Exercise 15.9.[2 ] Consider the source AS = {a, b, c, d, e}, PS = 1 1 1 1 1 / / / / / { 3, 3, 9, 9, 9} and the channel whose transition probability matrix is   1 0 0 0  0 0 2/3 0   Q= (15.1)  0 1 0 1 . 1 0 0 /3 0 Note that the source alphabet has five symbols, but the channel alphabet AX = AY = {0, 1, 2, 3} has only four. Assume that the source produces symbols at exactly 3/4 the rate that the channel accepts channel symbols. For a given (tiny)  > 0, explain how you would design a system for communicating the source’s output over the channel with an average error probability per source symbol less than . Be as explicit as possible. In particular, do not invoke Shannon’s noisy-channel coding theorem.

. Exercise 15.10. [2 ] Consider a binary symmetric channel and a code C = {0000, 0011, 1100, 1111}; assume that the four codewords are used with probabilities {1/2, 1/8, 1/8, 1/4}. What is the decoding rule that minimizes the probability of decoding error? [The optimal decoding rule depends on the noise level f of the binary symmetric channel. Give the decoding rule for each range of values of f , for f between 0 and 1/2.]

Exercise 15.11. [2 ] Find the capacity and optimal input distribution for the three-input, three-output channel whose transition probabilities are:   1 0 0 (15.2) Q =  0 2/3 1/3  . 0 1/3 2/3

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15 — Further Exercises on Information Theory

235

Exercise 15.12. [3, p.239] The input to a channel Q is a word of 8 bits. The output is also a word of 8 bits. Each time it is used, the channel flips exactly one of the transmitted bits, but the receiver does not know which one. The other seven bits are received without error. All 8 bits are equally likely to be the one that is flipped. Derive the capacity of this channel. Show, by describing an explicit encoder and decoder that it is possible reliably (that is, with zero error probability) to communicate 5 bits per cycle over this channel. . Exercise 15.13. [2 ] A channel with input x ∈ {a, b, c} and output y ∈ {r, s, t, u} has conditional probability matrix:  1  /2 0 0 * r  aH  1/2 1/2 0  j Hs  *   Q= bH  0 1/2 1/2  . j Ht *   cH 0 0 1/2 j Hu What is its capacity?

. Exercise 15.14. [3 ] The ten-digit number on the cover of a book known as the ISBN incorporates an error-detecting code. The number consists of nine source digits x1 , x2 , . . . , x9 , satisfying xn ∈ {0, 1, . . . , 9}, and a tenth check digit whose value is given by ! 9 X x10 = nxn mod 11. n=1

Here x10 ∈ {0, 1, . . . , 9, 10}. If x10 = 10 then the tenth digit is shown using the roman numeral X. Show that a valid ISBN satisfies: 10 X

n=1

nxn

!

mod 11 = 0.

Imagine that an ISBN is communicated over an unreliable human channel which sometimes modifies digits and sometimes reorders digits. Show that this code can be used to detect (but not correct) all errors in which any one of the ten digits is modified (for example, 1-010-00000-4 → 1-010-00080-4).

Show that this code can be used to detect all errors in which any two adjacent digits are transposed (for example, 1-010-00000-4 → 1-100-000004). What other transpositions of pairs of non-adjacent digits can be detected? If the tenth digit were defined to be x10 =

9 X

n=1

nxn

!

mod 10,

why would the code not work so well? (Discuss the detection of both modifications of single digits and transpositions of digits.)

0-521-64298-1 1-010-00000-4 Table 15.1. Some valid ISBNs. [The hyphens are included for legibility.]

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15 — Further Exercises on Information Theory

Exercise 15.15. [3 ] A channel with input x and output y has transition probability matrix:   1−f f 0 0  f 1−f 0 0  . Q=  0 0 1−g g  0 0 g 1−g

a - a @

R @ b - b

c - c @

R @ d - d

Assuming an input distribution of the form   p p 1−p 1−p PX = , , , , 2 2 2 2

write down the entropy of the output, H(Y ), and the conditional entropy of the output given the input, H(Y |X). Show that the optimal input distribution is given by p= where H2 (f ) = f log 2

1 f

1 , 1 + 2−H2 (g)+H2 (f )

+ (1 − f ) log 2

1 (1−f ) .

Remember

Write down the optimal input distribution and the capacity of the channel in the case f = 1/2, g = 0, and comment on your answer. . Exercise 15.16. [2 ] What are the differences in the redundancies needed in an error-detecting code (which can reliably detect that a block of data has been corrupted) and an error-correcting code (which can detect and correct errors)?

Further tales from information theory The following exercises give you the chance to discover for yourself the answers to some more surprising results of information theory. Exercise 15.17. [3 ] Communication of information from correlated sources. Imagine that we want to communicate data from two data sources X (A) and X (B) to a central location C via noise-free one-way communication channels (figure 15.2a). The signals x(A) and x(B) are strongly dependent, so their joint information content is only a little greater than the marginal information content of either of them. For example, C is a weather collator who wishes to receive a string of reports saying whether it is raining in Allerton (x (A) ) and whether it is raining in Bognor (x(B) ). The joint probability of x(A) and x(B) might be P (x(A) , x(B) ): x(A) 0 1 x(B)

0 1

0.49 0.01 0.01 0.49

(15.3)

The weather collator would like to know N successive values of x (A) and x(B) exactly, but, since he has to pay for every bit of information he receives, he is interested in the possibility of avoiding buying N bits from source A and N bits from source B. Assuming that variables x (A) and x(B) are generated repeatedly from this distribution, can they be encoded at rates R A and RB in such a way that C can reconstruct all the variables, with the sum of information transmission rates on the two lines being less than two bits per cycle?

d dp H2 (p)

= log2

1−p p .

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15 — Further Exercises on Information Theory

237 RB

x(A) (a)

x(B)

encode - t(A) H RA H j H C *   encode (B)  -t RB

H(X

(A)

,X

(B)

)

H(X (B) )

H(X (B) | X (A) )

(b)

             Achievable             

H(X (A) | X (B) ) H(X (A) )

RA

The answer, which you should demonstrate, is indicated in figure 15.2. In the general case of two dependent sources X (A) and X (B) , there exist codes for the two transmitters that can achieve reliable communication of both X (A) and X (B) to C, as long as: the information rate from X (A) , RA , exceeds H(X (A) | X (B) ); the information rate from X (B) , RB , exceeds H(X (B) | X (A) ); and the total information rate RA + RB exceeds the joint entropy H(X (A) , X (B) ) (Slepian and Wolf, 1973). So in the case of x(A) and x(B) above, each transmitter must transmit at a rate greater than H2 (0.02) = 0.14 bits, and the total rate R A + RB must be greater than 1.14 bits, for example R A = 0.6, RB = 0.6. There exist codes that can achieve these rates. Your task is to figure out why this is so. Try to find an explicit solution in which one of the sources is sent as plain text, t(B) = x(B) , and the other is encoded.

Figure 15.2. Communication of information from dependent sources. (a) x(A) and x(B) are dependent sources (the dependence is represented by the dotted arrow). Strings of values of each variable are encoded using codes of rate RA and RB into transmissions t(A) and t(B) , which are communicated over noise-free channels to a receiver C. (b) The achievable rate region. Both strings can be conveyed without error even though RA < H(X (A) ) and RB < H(X (B) ).

Exercise 15.18. [3 ] Multiple access channels. Consider a channel with two sets of inputs and one output – for example, a shared telephone line (figure 15.3a). A simple model system has two binary inputs x (A) and x(B) and a ternary output y equal to the arithmetic sum of the two inputs, that’s 0, 1 or 2. There is no noise. Users A and B cannot communicate with each other, and they cannot hear the output of the channel. If the output is a 0, the receiver can be certain that both inputs were set to 0; and if the output is a 2, the receiver can be certain that both inputs were set to 1. But if the output is 1, then it could be that the input state was (0, 1) or (1, 0). How should users A and B use this channel so that their messages can be deduced from the received signals? How fast can A and B communicate? Clearly the total information rate from A and B to the receiver cannot be two bits. On the other hand it is easy to achieve a total information rate RA +RB of one bit. Can reliable communication be achieved at rates (R A , RB ) such that RA + RB > 1? The answer is indicated in figure 15.3. Some practical codes for multi-user channels are presented in Ratzer and MacKay (2003). Exercise 15.19. [3 ] Broadcast channels. A broadcast channel consists of a single transmitter and two or more receivers. The properties of the channel are defined by a conditional distribution Q(y (A) , y (B) | x). (We’ll assume the channel is memoryless.) The task is to add an encoder and two decoders to enable reliable communication of a common message at rate R 0 to both receivers, an individual message at rate RA to receiver A, and an individual message at rate RB to receiver B. The capacity region of the broadcast channel is the convex hull of the set of achievable rate triplets (R 0 , RA , RB ). A simple benchmark for such a channel is given by time-sharing (timedivision signaling). If the capacities of the two channels, considered separately,

(A) *  y x Hj H (B) y

Figure 15.4. The broadcast channel. x is the channel input; y (A) and y (B) are the outputs.

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238

15 — Further Exercises on Information Theory

(a)

Figure 15.3. Multiple access channels. (a) A general multiple access channel with two transmitters and one receiver. (b) A binary multiple access channel with output equal to the sum of two inputs. (c) The achievable region.

-

x(A)

-

x(B)

- y

P (y|x(A) , x(B) )

RB 1

x(A) 0 1

y:

(b)

x(B)

0 1

0 1 1 2

1/2

Achievable

(c) 1/2

1 RA

are C (A) and C (B) , then by devoting a fraction φA of the transmission time to channel A and φB = 1−φA to channel B, we can achieve (R0 , RA , RB ) = (0, φA C (A) , φB C (B) ). We can do better than this, however. As an analogy, imagine speaking simultaneously to an American and a Belarusian; you are fluent in American and in Belarusian, but neither of your two receivers understands the other’s language. If each receiver can distinguish whether a word is in their own language or not, then an extra binary file can be conveyed to both recipients by using its bits to decide whether the next transmitted word should be from the American source text or from the Belarusian source text. Each recipient can concatenate the words that they understand in order to receive their personal message, and can also recover the binary string. An example of a broadcast channel consists of two binary symmetric channels with a common input. The two halves of the channel have flip probabilities fA and fB . We’ll assume that A has the better half-channel, i.e., fA < fB < 1/2. [A closely related channel is a ‘degraded’ broadcast channel, in which the conditional probabilities are such that the random variables have the structure of a Markov chain, x → y (A) → y (B) ,

RB C (B)

Exercise 15.20. [3 ] Variable-rate error-correcting codes for channels with unknown noise level. In real life, channels may sometimes not be well characterized before the encoder is installed. As a model of this situation, imagine that a channel is known to be a binary symmetric channel with noise level either f A or fB . Let fB > fA , and let the two capacities be CA and CB . Those who like to live dangerously might install a system designed for noise level fA with rate RA ' CA ; in the event that the noise level turns out to be fB , our experience of Shannon’s theories would lead us to expect that there

@

@

@ @ C (A) RA

Figure 15.5. Rates achievable by simple timesharing.

R

(15.4)

i.e., y (B) is a further degraded version of y (A) .] In this special case, it turns out that whatever information is getting through to receiver B can also be recovered by receiver A. So there is no point distinguishing between R 0 and RB : the task is to find the capacity region for the rate pair (R 0 , RA ), where R0 is the rate of information reaching both A and B, and R A is the rate of the extra information reaching A. The following exercise is equivalent to this one, and a solution to it is illustrated in figure 15.8.

6 @

CA CB fA

fB

f

Figure 15.6. Rate of reliable communication R, as a function of noise level f , for Shannonesque codes designed to operate at noise levels fA (solid line) and fB (dashed line).

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15 — Further Exercises on Information Theory would be a catastrophic failure to communicate information reliably (solid line in figure 15.6). A conservative approach would design the encoding system for the worstcase scenario, installing a code with rate R B ' CB (dashed line in figure 15.6). In the event that the lower noise level, f A , holds true, the managers would have a feeling of regret because of the wasted capacity difference C A − RB . Is it possible to create a system that not only transmits reliably at some rate R0 whatever the noise level, but also communicates some extra, ‘lowerpriority’ bits if the noise level is low, as shown in figure 15.7? This code communicates the high-priority bits reliably at all noise levels between f A and fB , and communicates the low-priority bits also if the noise level is f A or below. This problem is mathematically equivalent to the previous problem, the degraded broadcast channel. The lower rate of communication was there called R0 , and the rate at which the low-priority bits are communicated if the noise level is low was called RA . An illustrative answer is shown in figure 15.8, for the case f A = 0.01 and fB = 0.1. (This figure also shows the achievable region for a broadcast channel whose two half-channels have noise levels f A = 0.01 and fB = 0.1.) I admit I find the gap between the simple time-sharing solution and the cunning solution disappointingly small. In Chapter 50 we will discuss codes for a special class of broadcast channels, namely erasure channels, where every symbol is either received without error or erased. These codes have the nice property that they are rateless – the number of symbols transmitted is determined on the fly such that reliable comunication is achieved, whatever the erasure statistics of the channel. Exercise 15.21. [3 ] Multiterminal information networks are both important practically and intriguing theoretically. Consider the following example of a two-way binary channel (figure 15.9a,b): two people both wish to talk over the channel, and they both want to hear what the other person is saying; but you can hear the signal transmitted by the other person only if you are transmitting a zero. What simultaneous information rates from A to B and from B to A can be achieved, and how? Everyday examples of such networks include the VHF channels used by ships, and computer ethernet networks (in which all the devices are unable to hear anything if two or more devices are broadcasting simultaneously). Obviously, we can achieve rates of 1/2 in both directions by simple timesharing. But can the two information rates be made larger? Finding the capacity of a general two-way channel is still an open problem. However, we can obtain interesting results concerning achievable points for the simple binary channel discussed above, as indicated in figure 15.9c. There exist codes that can achieve rates up to the boundary shown. There may exist better codes too.

Solutions Solution to exercise 15.12 (p.235). C(Q) = 5 bits. Hint for the last part: a solution exists that involves a simple (8, 5) code.

239

R CA CB f

f

A

B

f

Figure 15.7. Rate of reliable communication R, as a function of noise level f , for a desired variable-rate code.

0.6 0.4 0.2 0 0

0.2

0.4

0.6

0.8

1

Figure 15.8. An achievable region for the channel with unknown noise level. Assuming the two possible noise levels are fA = 0.01 and fB = 0.1, the dashed lines show the rates RA , RB that are achievable using a simple time-sharing approach, and the solid line shows rates achievable using a more cunning approach.

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240

15 — Further Exercises on Information Theory

-

x(A)

P (y (A) , y (B) |x(A) , x(B) )



y (A) :

(b)

x(B)

x(A) 0 1



x(B)

y (B) :

0 0 1 0

0 1

x(B)

x(A) 0 1 0 1

0 1 0 0

0.8

1

1

0.8

0.6

Achievable

R(B)

(a)

y (A)

- y (B)

0.4

0.2

(c)

0 0

0.2

0.4 0.6 R(A)

Figure 15.9. (a) A general two-way channel. (b) The rules for a binary two-way channel. The two tables show the outputs y (A) and y (B) that result for each state of the inputs. (c) Achievable region for the two-way binary channel. Rates below the solid line are achievable. The dotted line shows the ‘obviously achievable’ region which can be attained by simple time-sharing.

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16 Message Passing One of the themes of this book is the idea of doing complicated calculations using simple distributed hardware. It turns out that quite a few interesting problems can be solved by message-passing algorithms, in which simple messages are passed locally among simple processors whose operations lead, after some time, to the solution of a global problem.

16.1 Counting As an example, consider a line of soldiers walking in the mist. The commander wishes to perform the complex calculation of counting the number of soldiers in the line. This problem could be solved in two ways. First there is a solution that uses expensive hardware: the loud booming voices of the commander and his men. The commander could shout ‘all soldiers report back to me within one minute!’, then he could listen carefully as the men respond ‘Molesworth here sir!’, ‘Fotherington–Thomas here sir!’, and so on. This solution relies on several expensive pieces of hardware: there must be a reliable communication channel to and from every soldier; the commander must be able to listen to all the incoming messages – even when there are hundreds of soldiers – and must be able to count; and all the soldiers must be well-fed if they are to be able to shout back across the possibly-large distance separating them from the commander. The second way of finding this global function, the number of soldiers, does not require global communication hardware, high IQ, or good food; we simply require that each soldier can communicate single integers with the two adjacent soldiers in the line, and that the soldiers are capable of adding one to a number. Each soldier follows these rules: 1. If you are the front soldier in the line, say the number ‘one’ to the soldier behind you. 2. If you are the rearmost soldier in the line, say the number ‘one’ to the soldier in front of you. 3. If a soldier ahead of or behind you says a number to you, add one to it, and say the new number to the soldier on the other side.

If the clever commander can not only add one to a number, but also add two numbers together, then he can find the global number of soldiers by simply adding together:

241

Algorithm 16.1. Message-passing rule-set A.

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242

+

+

16 — Message Passing

the number said to him by the soldier in front of him, the number said to the commander by the soldier behind him, one

(which equals the total number of soldiers in front) (which is the number behind)

(to count the commander himself).

This solution requires only local communication hardware and simple computations (storage and addition of integers). 1

2

3

4

3

2

4

Commander

1

Figure 16.2. A line of soldiers counting themselves using message-passing rule-set A. The commander can add ‘3’ from the soldier in front, ‘1’ from the soldier behind, and ‘1’ for himself, and deduce that there are 5 soldiers in total.

Separation This clever trick makes use of a profound property of the total number of soldiers: that it can be written as the sum of the number of soldiers in front of a point and the number behind that point, two quantities which can be computed separately, because the two groups are separated by the commander. If the soldiers were not arranged in a line but were travelling in a swarm, then it would not be easy to separate them into two groups in this way. The Figure 16.3. A swarm of guerillas.

Commander

Jim

guerillas in figure 16.3 could not be counted using the above message-passing rule-set A, because, while the guerillas do have neighbours (shown by lines), it is not clear who is ‘in front’ and who is ‘behind’; furthermore, since the graph of connections between the guerillas contains cycles, it is not possible for a guerilla in a cycle (such as ‘Jim’) to separate the group into two groups, ‘those in front’, and ‘those behind’. A swarm of guerillas can be counted by a modified message-passing algorithm if they are arranged in a graph that contains no cycles. Rule-set B is a message-passing algorithm for counting a swarm of guerillas whose connections form a cycle-free graph, also known as a tree, as illustrated in figure 16.4. Any guerilla can deduce the total in the tree from the messages that they receive.

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16.1: Counting

243

Figure 16.4. A swarm of guerillas whose connections form a tree.

Commander

Jim

Algorithm 16.5. Message-passing rule-set B.

1. Count your number of neighbours, N . 2. Keep count of the number of messages you have received from your neighbours, m, and of the values v1 , v2 , . . . , vN of each of those messages. Let V be the running total of the messages you have received. 3. If the number of messages you have received, m, is equal to N − 1, then identify the neighbour who has not sent you a message and tell them the number V + 1. 4. If the number of messages you have received is equal to N , then: (a) the number V + 1 is the required total. (b) for each neighbour n { say to neighbour n the number V + 1 − vn . }

A

Figure 16.6. A triangular 41 × 41 grid. How many paths are there from A to B? One path is shown.

B

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244

16 — Message Passing

16.2 Path-counting A more profound task than counting squaddies is the task of counting the number of paths through a grid, and finding how many paths pass through any given point in the grid. Figure 16.6 shows a rectangular grid, and a path through the grid, connecting points A and B. A valid path is one that starts from A and proceeds to B by rightward and downward moves. Our questions are:

A N P M

1. How many such paths are there from A to B? 2. If a random path from A to B is selected, what is the probability that it passes through a particular node in the grid? [When we say ‘random’, we mean that all paths have exactly the same probability of being selected.] 3. How can a random path from A to B be selected? Counting all the paths from A to B doesn’t seem straightforward. The number of paths is expected to be pretty big – even if the permitted grid were a diagonal strip only three nodes wide, there would still be about 2 N/2 possible paths. The computational breakthrough is to realize that to find the number of paths, we do not have to enumerate all the paths explicitly. Pick a point P in the grid and consider the number of paths from A to P. Every path from A to P must come in to P through one of its upstream neighbours (‘upstream’ meaning above or to the left). So the number of paths from A to P can be found by adding up the number of paths from A to each of those neighbours. This message-passing algorithm is illustrated in figure 16.8 for a simple grid with ten vertices connected by twelve directed edges. We start by sending the ‘1’ message from A. When any node has received messages from all its upstream neighbours, it sends the sum of them on to its downstream neighbours. At B, the number 5 emerges: we have counted the number of paths from A to B without enumerating them all. As a sanity-check, figure 16.9 shows the five distinct paths from A to B. Having counted all paths, we can now move on to more challenging problems: computing the probability that a random path goes through a given vertex, and creating a random path.

B Figure 16.7. Every path from A to P enters P through an upstream neighbour of P, either M or N; so we can find the number of paths from A to P by adding the number of paths from A to M to the number from A to N. 1

A

1

1

1

1

2

3

2

5 5

B

Figure 16.8. Messages sent in the forward pass.

A

B

Probability of passing through a node By making a backward pass as well as the forward pass, we can deduce how many of the paths go through each node; and if we divide that by the total number of paths, we obtain the probability that a randomly selected path passes through that node. Figure 16.10 shows the backward-passing messages in the lower-right corners of the tables, and the original forward-passing messages in the upper-left corners. By multiplying these two numbers at a given vertex, we find the total number of paths passing through that vertex. For example, four paths pass through the central vertex. Figure 16.11 shows the result of this computation for the triangular 41 × 41 grid. The area of each blob is proportional to the probability of passing through the corresponding node.

Random path sampling Exercise 16.1.[1, p.247] If one creates a ‘random’ path from A to B by flipping a fair coin at every junction where there is a choice of two directions, is

Figure 16.9. The five paths. 1

1

A

1 5

5

1 3

2

1

1 3

2

2 2

1 5

1

B

1 5 1

Figure 16.10. Messages sent in the forward and backward passes.

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16.3: Finding the lowest-cost path

245

the resulting path a uniform random sample from the set of all paths? [Hint: imagine trying it for the grid of figure 16.8.] There is a neat insight to be had here, and I’d like you to have the satisfaction of figuring it out. Exercise 16.2.[2, p.247] Having run the forward and backward algorithms between points A and B on a grid, how can one draw one path from A to B uniformly at random? (Figure 16.11.) A

(a)

(b)

Figure 16.11. (a) The probability of passing through each node, and (b) a randomly chosen path.

B

The message-passing algorithm we used to count the paths to B is an example of the sum–product algorithm. The ‘sum’ takes place at each node when it adds together the messages coming from its predecessors; the ‘product’ was not mentioned, but you can think of the sum as a weighted sum in which all the summed terms happened to have weight 1.

16.3 Finding the lowest-cost path Imagine you wish to travel as quickly as possible from Ambridge (A) to Bognor (B). The various possible routes are shown in figure 16.12, along with the cost in hours of traversing each edge in the graph. For example, the route A–I–L– N–B has a cost of 8 hours. We would like to find the lowest-cost path without explicitly evaluating the cost of all paths. We can do this efficiently by finding for each node what the cost of the lowest-cost path to that node from A is. These quantities can be computed by message-passing, starting from node A. The message-passing algorithm is called the min–sum algorithm or Viterbi algorithm. For brevity, we’ll call the cost of the lowest-cost path from node A to node x ‘the cost of x’. Each node can broadcast its cost to its descendants once it knows the costs of all its possible predecessors. Let’s step through the algorithm by hand. The cost of A is zero. We pass this news on to H and I. As the message passes along each edge in the graph, the cost of that edge is added. We find the costs of H and I are 4 and 1 respectively (figure 16.13a). Similarly then, the costs of J and L are found to be 6 and 2 respectively, but what about K? Out of the edge H–K comes the message that a path of cost 5 exists from A to K via H; and from edge I–K we learn of an alternative path of cost 3 (figure 16.13b). The min–sum algorithm sets the cost of K equal to the minimum of these (the ‘min’), and records which was the smallest-cost route into K by retaining only the edge I–K and pruning away the other edges leading to K (figure 16.13c). Figures 16.13d and e show the remaining two iterations of the algorithm which reveal that there is a path from A to B with cost 6. [If the min–sum algorithm encounters a tie, where the minimum-cost

J H2 2* H  j H   H M H1 1 H 4* 2* H H   j H j H     AH KH B 2 H H *  *  j H j H  1 1 N 3 I H H *  j H 1 L 3

Figure 16.12. Route diagram from Ambridge to Bognor, showing the costs associated with the edges.

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246

16 — Message Passing

path to a node is achieved by more than one route to it, then the algorithm can pick any of those routes at random.] We can recover this lowest-cost path by backtracking from B, following the trail of surviving edges back to A. We deduce that the lowest-cost path is A–I–K–M–B. 2* J H2H j 4 1 M H1 2 4 HH H H * * j j  0  KH B 2 AH H H * * j j  1 N 3 1 1

(a)

Other applications of the min–sum algorithm Imagine that you manage the production of a product from raw materials via a large set of operations. You wish to identify the critical path in your process, that is, the subset of operations that are holding up production. If any operations on the critical path were carried out a little faster then the time to get from raw materials to product would be reduced. The critical path of a set of operations can be found using the min–sum algorithm. In Chapter 25 the min–sum algorithm will be used in the decoding of error-correcting codes.

H

I

H j

1

* 

L 3

6 2* J H2H j 4 1 M H1 2 4 H Hj H H * * j 5  0 K B 2 3 HH AH H * j j  *  1 N 3 1 1 H I H * j  1 2 3

(b)

L

16.4 Summary and related ideas

(c)

Some global functions have a separability property. For example, the number of paths from A to P separates into the sum of the number of paths from A to M (the point to P’s left) and the number of paths from A to N (the point above P). Such functions can be computed efficiently by message-passing. Other functions do not have such separability properties, for example

4* 0  A

3. the length of the shortest tour that a travelling salesman could take that visits every soldier in a troop. One of the challenges of machine learning is to find low-cost solutions to problems like these. The problem of finding a large subset of variables that are approximately equal can be solved with a neural network approach (Hopfield and Brody, 2000; Hopfield and Brody, 2001). A neural approach to the travelling salesman problem will be discussed in section 42.9.

16.5 Further exercises . Exercise 16.3.[2 ] Describe the asymptotic properties of the probabilities depicted in figure 16.11a, for a grid in a triangle of width and height N .

H

6 J

H2 H j

2*M H1H j 3  B

H H * *K H j j 1 2  1 N 3  H I H * j  1 2 3

1

L

(d)

1. the number of pairs of soldiers in a troop who share the same birthday; 2. the size of the largest group of soldiers who share a common height (rounded to the nearest centimetre);

H

2* 4 1

4* 

0 A

H

2* 4 1

6 J

2 5 1 H 2* M H j

H

3 B H H * *K H j j 1 2  1 4 3  N I H H j 1 2 3

1

L

(e) 4* 0  A

2* 4 1

H

6 J

2

5 2* M H1H j6 3 

HH 2* K HH j j 1 1  1 4 3 H I N H j 1 2 3 L

Figure 16.13. Min–sum message-passing algorithm to find the cost of getting to each node, and thence the lowest cost route from A to B.

. Exercise 16.4.[2 ] In image processing, the integral image I(x, y) obtained from an image f (x, y) (where x and y are pixel coordinates) is defined by I(x, y) ≡

y x X X

f (u, v).

(16.1)

u=0 v=0

Show that the integral image I(x, y) can be efficiently computed by message passing. Show that, from the integral image, some simple functions of the image can be obtained. For example, give an expression for the sum of the image intensities f (x, y) for all (x, y) in a rectangular region extending from (x1 , y1 ) to (x2 , y2 ).

B

y2 y1 x1 (0, 0)

x2

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16.6: Solutions

16.6 Solutions Solution to exercise 16.1 (p.244). Since there are five paths through the grid of figure 16.8, they must all have probability 1/5. But a strategy based on fair coin-flips will produce paths whose probabilities are powers of 1/2. Solution to exercise 16.2 (p.245). To make a uniform random walk, each forward step of the walk should be chosen using a different biased coin at each junction, with the biases chosen in proportion to the backward messages emanating from the two options. For example, at the first choice after leaving A, there is a ‘3’ message coming from the East, and a ‘2’ coming from South, so one should go East with probability 3/5 and South with probability 2/5. This is how the path in figure 16.11b was generated.

247

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17 Communication over Constrained Noiseless Channels In this chapter we study the task of communicating efficiently over a constrained noiseless channel – a constrained channel over which not all strings from the input alphabet may be transmitted. We make use of the idea introduced in Chapter 16, that global properties of graphs can be computed by a local message-passing algorithm.

17.1 Three examples of constrained binary channels A constrained channel can be defined by rules that define which strings are permitted. Example 17.1. In Channel A every 1 must be followed by at least one 0. A valid string for this channel is 00100101001010100010.

Channel A: the substring 11 is forbidden.

(17.1)

As a motivation for this model, consider a channel in which 1s are represented by pulses of electromagnetic energy, and the device that produces those pulses requires a recovery time of one clock cycle after generating a pulse before it can generate another. Example 17.2. Channel B has the rule that all 1s must come in groups of two or more, and all 0s must come in groups of two or more. A valid string for this channel is 00111001110011000011.

Channel B: 101 and 010 are forbidden.

(17.2)

As a motivation for this model, consider a disk drive in which successive bits are written onto neighbouring points in a track along the disk surface; the values 0 and 1 are represented by two opposite magnetic orientations. The strings 101 and 010 are forbidden because a single isolated magnetic domain surrounded by domains having the opposite orientation is unstable, so that 101 might turn into 111, for example. Example 17.3. Channel C has the rule that the largest permitted runlength is two, that is, each symbol can be repeated at most once. A valid string for this channel is 10010011011001101001. 248

(17.3)

Channel C: 111 and 000 are forbidden.

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17.1: Three examples of constrained binary channels

249

A physical motivation for this model is a disk drive in which the rate of rotation of the disk is not known accurately, so it is difficult to distinguish between a string of two 1s and a string of three 1s, which are represented by oriented magnetizations of duration 2τ and 3τ respectively, where τ is the (poorly known) time taken for one bit to pass by; to avoid the possibility of confusion, and the resulting loss of synchronization of sender and receiver, we forbid the string of three 1s and the string of three 0s. All three of these channels are examples of runlength-limited channels. The rules constrain the minimum and maximum numbers of successive 1s and 0s. Channel unconstrained

A B C

Runlength of 1s Runlength of 0s minimum maximum minimum maximum ∞ 1 ∞ 2

1 1 2 1

1 1 2 1

∞ ∞ ∞ 2

In channel A, runs of 0s may be of any length but runs of 1s are restricted to length one. In channel B all runs must be of length two or more. In channel C, all runs must be of length one or two. The capacity of the unconstrained binary channel is one bit per channel use. What are the capacities of the three constrained channels? [To be fair, we haven’t defined the ‘capacity’ of such channels yet; please understand ‘capacity’ as meaning how many bits can be conveyed reliably per channel-use.]

Some codes for a constrained channel Let us concentrate for a moment on channel A, in which runs of 0s may be of any length but runs of 1s are restricted to length one. We would like to communicate a random binary file over this channel as efficiently as possible. A simple starting point is a (2, 1) code that maps each source bit into two transmitted bits, C1 . This is a rate-1/2 code, and it respects the constraints of channel A, so the capacity of channel A is at least 0.5. Can we do better? C1 is redundant because if the first of two received bits is a zero, we know that the second bit will also be a zero. We can achieve a smaller average transmitted length using a code that omits the redundant zeroes in C 1 . C2 is such a variable-length code. If the source symbols are used with equal frequency then the average transmitted length per source bit is 1 1 3 L= 1+ 2= , 2 2 2

(17.4)

so the average communication rate is R = 2/3,

(17.5)

and the capacity of channel A must be at least 2/3. Can we do better than C2 ? There are two ways to argue that the information rate could be increased above R = 2/3. The first argument assumes we are comfortable with the entropy as a measure of information content. The idea is that, starting from code C 2 , we can reduce the average message length, without greatly reducing the entropy

Code C1 s

t

0 00 1 10 Code C2 s

t

0 0 1 10

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250

17 — Communication over Constrained Noiseless Channels

of the message we send, by decreasing the fraction of 1s that we transmit. Imagine feeding into C2 a stream of bits in which the frequency of 1s is f . [Such a stream could be obtained from an arbitrary binary file by passing the source file into the decoder of an arithmetic code that is optimal for compressing binary strings of density f .] The information rate R achieved is the entropy of the source, H2 (f ), divided by the mean transmitted length, L(f ) = (1 − f ) + 2f = 1 + f. Thus

2

1+f H_2(f) 1

(17.6) 0 0

H2 (f ) H2 (f ) R(f ) = = . L(f ) 1+f

(17.7)

The original code C2 , without preprocessor, corresponds to f = 1/2. What happens if we perturb f a little towards smaller f , setting 1 f = + δ, 2

(17.8)

for small negative δ? In the vicinity of f = 1/2, the denominator L(f ) varies linearly with δ. In contrast, the numerator H 2 (f ) only has a second-order dependence on δ. . Exercise 17.4.[1 ] Find, to order δ 2 , the Taylor expansion of H2 (f ) as a function of δ. To first order, R(f ) increases linearly with decreasing δ. It must be possible to increase R by decreasing f . Figure 17.1 shows these functions; R(f ) does indeed increase as f decreases and has a maximum of about 0.69 bits per channel use at f ' 0.38. By this argument we have shown that the capacity of channel A is at least maxf R(f ) = 0.69. . Exercise 17.5.[2, p.257] If a file containing a fraction f = 0.5 1s is transmitted by C2 , what fraction of the transmitted stream is 1s? What fraction of the transmitted bits is 1s if we drive code C 2 with a sparse source of density f = 0.38? A second, more fundamental approach counts how many valid sequences of length N there are, SN . We can communicate log SN bits in N channel cycles by giving one name to each of these valid sequences.

17.2 The capacity of a constrained noiseless channel We defined the capacity of a noisy channel in terms of the mutual information between its input and its output, then we proved that this number, the capacity, was related to the number of distinguishable messages S(N ) that could be reliably conveyed over the channel in N uses of the channel by C = lim

N →∞

1 log S(N ). N

(17.9)

In the case of the constrained noiseless channel, we can adopt this identity as our definition of the channel’s capacity. However, the name s, which, when we were making codes for noisy channels (section 9.6), ran over messages s = 1, . . . , S, is about to take on a new role: labelling the states of our channel;

0.25

0.5

0.75

1

0.75

1

0.7 0.6 0.5 0.4 0.3

R(f) = H_2(f)/(1+f)

0.2 0.1 0 0

0.25

0.5

Figure 17.1. Top: The information content per source symbol and mean transmitted length per source symbol as a function of the source density. Bottom: The information content per transmitted symbol, in bits, as a function of f .

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17.3: Counting the number of possible messages s1 s2 s3 s4 s5 s6 s7 s8 f 1f 1f 1f 1f 1f 1f 1f 1 01 01 01 01 01 01 01 1  @  @  @  @  @  @  @  @@ @@ @@ @@ @@ @@ @@ - f R - f R - f R - f R - f R - f R - f R - f 0f 0 0 0 0 0 0 0 0 0 0 0 0 0 0 0 0

1 1 (a) 0

0 0 (c)

sn+1 sn j 1j 1 @0 1 @ @ R @ 0j 0 - 0j

(b)

1

11

1

0 1 0

1 0

0 00

B

sn m 11

251

(from)

(d)

sn+1 1 - 11 m 1

A A0 A 1m A

A   A A A  U 0m 1 0 @  @  @  R @ m 00 0  1 1 0  0 0 0 A=  1 0 0 0 0 1

1m 0m m 00

 0 1   0  1

A=

(to)

1 0

0  1 1

Figure 17.2. (a) State diagram for channel A. (b) Trellis section. (c) Trellis. (d) Connection matrix. sn+1

sn n 11

1 0 A A 1n A @0 A1   @ A  @AU A @ R 0n 1 0 @  @  @  R @ n 00

11

A

1 1 0 10 1 0 0 00

C

1 0 1



0  0 A=  1 0

1 0 1 0

0 1 0 1

n 11 1n 0n n 00  0 1   0  0

so in this chapter we will denote the number of distinguishable messages of length N by MN , and define the capacity to be: C = lim

N →∞

1 log MN . N

(17.10)

Once we have figured out the capacity of a channel we will return to the task of making a practical code for that channel.

17.3 Counting the number of possible messages First let us introduce some representations of constrained channels. In a state diagram, states of the transmitter are represented by circles labelled with the name of the state. Directed edges from one state to another indicate that the transmitter is permitted to move from the first state to the second, and a label on that edge indicates the symbol emitted when that transition is made. Figure 17.2a shows the state diagram for channel A. It has two states, 0 and 1. When transitions to state 0 are made, a 0 is transmitted; when transitions to state 1 are made, a 1 is transmitted; transitions from state 1 to state 1 are not possible. We can also represent the state diagram by a trellis section, which shows two successive states in time at two successive horizontal locations (figure 17.2b). The state of the transmitter at time n is called s n . The set of possible state sequences can be represented by a trellis as shown in figure 17.2c. A valid sequence corresponds to a path through the trellis, and the number of

Figure 17.3. State diagrams, trellis sections and connection matrices for channels B and C.

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252

17 — Communication over Constrained Noiseless Channels

M1 = 2

 h

0

1 1h

- 0h 1

h

0

(b) Channel B

1 1h  @

h

- 0h

0

@



1 1h

@ R - 0h

M2 = 3

M1 = 2

h

0

2

1

M1 = 2

(a) Channel A

M2 = 3

M1 = 2

M3 = 5

M2 = 3

Figure 17.4. Counting the number of paths in the trellis of channel A. The counts next to the nodes are accumulated by passing from left to right across the trellises.

M3 = 5

1 2 1 1h 1h 1h  @  @  @ @ R @ R @ - 0h - 0h - 0h 2

1

3

M4 = 8 M5 = 13 M6 = 21 M7 = 34 M8 = 55

1 2 3 5 8 13 21 1h 1h 1h 1h 1h 1h 1h  @  @  @  @  @  @  @ @ @ @ @ @ @ @ R R @ R @ @ R R @ R @ R @ - 0h - 0h - 0h - 0h - 0h - 0h - 0h - 0h

1 1h  @ 1

2

3

M1 = 2

M2 = 3

M3 = 5

5

8

13

21

34

M4 = 8 M5 = 13 M6 = 21 M7 = 34 M8 = 55

1 2 3 4 6 10 17 - 11h - 11h - 11h - 11h - 11h - 11h - 11h h  A  A  A  A  A  A  A 1 A 1 A 1 A 2 A 4 A 7 A 11 1 A 1h A 1h A 1h A 1h A 1h A 1h 1h A 1h A A  A  A  A  A  A  A         A A A A A A A    1  2  3   6  10 A h  U A h  U A h  U A h  AU 4h  U A h  U A h  0h  U 0 0 0 0 0 0 0  @ @ @ @ @ @ @  @ @ @ @ @ @ @  - h @  R  R  @  R  R  R R @ @ R @ @ @ - 00h h h h h h h h 00 00 00 00 00 00 00 00 11

(c) Channel C

1

1

1

2

M1 = 1

M2 = 2

M3 = 3

M4 = 5

4

7

11

17

M5 = 8 M6 = 13 M7 = 21 M8 = 34

7 4 2 2 1 1 h h h h h h h h 11 11 11 11 11 11 11  A  A  A   A  A  A A A 10 7 A 4 A 2 A 2 A 1 A 1 A 1h 1h A 1h A 1h A 1h A 1h A 1h A 1h A       @ A  @ A  @ A  @ A  @ A  @ A  @ A @A  11  6 @A  4 @A  3 @A  1 @A  1 @A  1 @A  A h A h R A U U A U A U A U A @U R @ R @ @ R R @ R @ @ R h h h h h h  U  0  0  0  0  0  0 0 0  @ @ @ @ @ @ @ @ @ @ @ @ @ @   R  R  R  @  R  R  @  @ 00h @ @ @ @ R R h h h h h h h h 00 00 00 00 00 00 00 00 11

1

1

1

3

4

6

Figure 17.5. Counting the number of paths in the trellises of channels A, B, and C. We assume that at the start the first bit is preceded by 00, so that for channels A and B, any initial character is permitted, but for channel C, the first character must be a 1.

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17.3: Counting the number of possible messages

253

n

Mn

Mn /Mn−1

log2 Mn

1 2 3 4 5 6 7 8 9 10 11 12 100 200 300 400

2 3 5 8 13 21 34 55 89 144 233 377 9×1020 7×1041 6×1062 5×1083

1.500 1.667 1.600 1.625 1.615 1.619 1.618 1.618 1.618 1.618 1.618 1.618 1.618 1.618 1.618

1.0 1.6 2.3 3.0 3.7 4.4 5.1 5.8 6.5 7.2 7.9 8.6 69.7 139.1 208.5 277.9

1 n

Figure 17.6. Counting the number of paths in the trellis of channel A.

log 2 Mn 1.00 0.79 0.77 0.75 0.74 0.73 0.73 0.72 0.72 0.72 0.71 0.71 0.70 0.70 0.70 0.69

valid sequences is the number of paths. For the purpose of counting how many paths there are through the trellis, we can ignore the labels on the edges and summarize the trellis section by the connection matrix A, in which A ss0 = 1 if there is an edge from state s to s0 , and Ass0 = 0 otherwise (figure 17.2d). Figure 17.3 shows the state diagrams, trellis sections and connection matrices for channels B and C. Let’s count the number of paths for channel A by message-passing in its trellis. Figure 17.4 shows the first few steps of this counting process, and figure 17.5a shows the number of paths ending in each state after n steps for n = 1, . . . , 8. The total number of paths of length n, M n , is shown along the top. We recognize Mn as the Fibonacci series. . Exercise 17.6.[1 ] Show that the ratio of successive terms in the Fibonacci series tends to the golden ratio, √ 1+ 5 = 1.618. (17.11) γ≡ 2 Thus, to within a constant factor, M N scales as MN ∼ γ N as N → ∞, so the capacity of channel A is   1 (17.12) C = lim log 2 constant · γ N = log 2 γ = log2 1.618 = 0.694. N How can we describe what we just did? The count of the number of paths is a vector c(n) ; we can obtain c(n+1) from c(n) using: c(n+1) = Ac(n) .

(17.13)

c(N ) = AN c(0) ,

(17.14)

So c(0)

where is the state count before any symbols are transmitted. In figure 17.5 we assumed c(0) = [0, 1]T, i.e., that either of the two symbols is permitted at P (n) the outset. The total number of paths is M n = s cs = c(n) · n. In the limit, c(N ) becomes dominated by the principal right-eigenvector of A. (0)

c(N ) → constant · λN 1 eR .

(17.15)

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254

17 — Communication over Constrained Noiseless Channels

Here, λ1 is the principal eigenvalue of A. So to find the capacity of any constrained channel, all we need to do is find the principal eigenvalue, λ1 , of its connection matrix. Then C = log2 λ1 .

(17.16)

17.4 Back to our model channels Comparing figure 17.5a and figures 17.5b and c it looks as if channels B and C have the same capacity as channel A. The principal eigenvalues of the three trellises are the same (the eigenvectors for channels A and B are given at the bottom of table C.4, p.608). And indeed the channels are intimately related.

-

t

z1 hd z0

6 -⊕

s

z1 hd z0 

t

? -⊕ - s

Figure 17.7. An accumulator and a differentiator.

Equivalence of channels A and B If we take any valid string s for channel A and pass it through an accumulator, obtaining t defined by: t1 = s 1 tn = tn−1 + sn mod 2 for n ≥ 2,

(17.17)

then the resulting string is a valid string for channel B, because there are no 11s in s, so there are no isolated digits in t. The accumulator is an invertible operator, so, similarly, any valid string t for channel B can be mapped onto a valid string s for channel A through the binary differentiator, s1 = t 1 sn = tn − tn−1 mod 2 for n ≥ 2.

(17.18)

Because + and − are equivalent in modulo 2 arithmetic, the differentiator is also a blurrer, convolving the source stream with the filter (1, 1). Channel C is also intimately related to channels A and B. . Exercise 17.7.[1, p.257] What is the relationship of channel C to channels A and B?

17.5 Practical communication over constrained channels OK, how to do it in practice? Since all three channels are equivalent, we can concentrate on channel A.

s

c(s)

1 2 3 4 5 6 7 8

00000 10000 01000 00100 00010 10100 01010 10010

Fixed-length solutions We start with explicitly-enumerated codes. The code in the table 17.8 achieves a rate of 3/5 = 0.6. . Exercise 17.8.[1, p.257] Similarly, enumerate all strings of length 8 that end in the zero state. (There are 34 of them.) Hence show that we can map 5 bits (32 source strings) to 8 transmitted bits and achieve rate 5/8 = 0.625. What rate can be achieved by mapping an integer number of source bits to N = 16 transmitted bits?

Table 17.8. A runlength-limited code for channel A.

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17.5: Practical communication over constrained channels

255

Optimal variable-length solution The optimal way to convey information over the constrained channel is to find the optimal transition probabilities for all points in the trellis, Q s0 |s , and make transitions with these probabilities. When discussing channel A, we showed that a sparse source with density f = 0.38, driving code C2 , would achieve capacity. And we know how to make sparsifiers (Chapter 6): we design an arithmetic code that is optimal for compressing a sparse source; then its associated decoder gives an optimal mapping from dense (i.e., random binary) strings to sparse strings. The task of finding the optimal probabilities is given as an exercise. Exercise 17.9.[3 ] Show that the optimal transition probabilities Q can be found as follows. Find the principal right- and left-eigenvectors of A, that is the solutions T T of Ae(R) = λe(R) and e(L) A = λe(L) with largest eigenvalue λ. Then construct a matrix Q whose invariant distribution is proportional to (R) (L) ei ei , namely (L)

Qs0 |s =

e s0 A s0 s (L)

λes

.

(17.19)

[Hint: exercise 16.2 (p.245) might give helpful cross-fertilization here.] . Exercise 17.10. [3, p.258] Show that when sequences are generated using the optimal transition probability matrix (17.19), the entropy of the resulting sequence is asymptotically log 2 λ per symbol. [Hint: consider the conditional entropy of just one symbol given the previous one, assuming the previous one’s distribution is the invariant distribution.] In practice, we would probably use finite-precision approximations to the optimal variable-length solution. One might dislike variable-length solutions because of the resulting unpredictability of the actual encoded length in any particular case. Perhaps in some applications we would like a guarantee that the encoded length of a source file of size N bits will be less than a given length such as N/(C + ). For example, a disk drive is easier to control if all blocks of 512 bytes are known to take exactly the same amount of disk real-estate. For some constrained channels we can make a simple modification to our variable-length encoding and offer such a guarantee, as follows. We find two codes, two mappings of binary strings to variable-length encodings, having the property that for any source string x, if the encoding of x under the first code is shorter than average, then the encoding of x under the second code is longer than average, and vice versa. Then to transmit a string x we encode the whole string with both codes and send whichever encoding has the shortest length, prepended by a suitably encoded single bit to convey which of the two codes is being used. [3C, p.258]

. Exercise 17.11. How many valid sequences of length 8 starting with a 0 are there for the run-length-limited channels shown in figure 17.9? What are the capacities of these channels? Using a computer, find the matrices Q for generating a random path through the trellises of the channel A, and the two run-length-limited channels shown in figure 17.9.

2 1





0 1 0  0 0 1  1 1 1 

0  0   0 1

1 0 0 1

0 1 0 1

0

1 1 0 0 0 3



0 0   1  1

0

1 2 1 1

0

1 0 0

0

Figure 17.9. State diagrams and connection matrices for channels with maximum runlengths for 1s equal to 2 and 3.

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256

17 — Communication over Constrained Noiseless Channels

. Exercise 17.12. [3, p.258] Consider the run-length-limited channel in which any length of run of 0s is permitted, and the maximum run length of 1s is a large number L such as nine or ninety. Estimate the capacity of this channel. (Give the first two terms in a series expansion involving L.) What, roughly, is the form of the optimal matrix Q for generating a random path through the trellis of this channel? Focus on the values of the elements Q1|0 , the probability of generating a 1 given a preceding 0, and QL|L−1 , the probability of generating a 1 given a preceding run of L−1 1s. Check your answer by explicit computation for the channel in which the maximum runlength of 1s is nine.

17.6 Variable symbol durations We can add a further frill to the task of communicating over constrained channels by assuming that the symbols we send have different durations, and that our aim is to communicate at the maximum possible rate per unit time. Such channels can come in two flavours: unconstrained, and constrained.

Unconstrained channels with variable symbol durations We encountered an unconstrained noiseless channel with variable symbol durations in exercise 6.18 (p.125). Solve that problem, and you’ve done this topic. The task is to determine the optimal frequencies with which the symbols should be used, given their durations. There is a nice analogy between this task and the task of designing an optimal symbol code (Chapter 4). When we make an binary symbol code for a source with unequal probabilities p i , the optimal message lengths are li∗ = log2 1/pi , so ∗ (17.20) pi = 2−li . Similarly, when we have a channel whose symbols have durations l i (in some units of time), the optimal probability with which those symbols should be used is (17.21) p∗i = 2−βli , where β is the capacity of the channel in bits per unit time.

Constrained channels with variable symbol durations Once you have grasped the preceding topics in this chapter, you should be able to figure out how to define and find the capacity of these, the trickiest constrained channels. Exercise 17.13. [3 ] A classic example of a constrained channel with variable symbol durations is the ‘Morse’ channel, whose symbols are the the the the

dot dash short space (used between letters in morse code) long space (used between words)

d, D, s, and S;

the constraints are that spaces may only be followed by dots and dashes. Find the capacity of this channel in bits per unit time assuming (a) that all four symbols have equal durations; or (b) that the symbol durations are 2, 4, 3 and 6 time units respectively.

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17.7: Solutions

257

Exercise 17.14. [4 ] How well-designed is Morse code for English (with, say, the probability distribution of figure 2.1)? Exercise 17.15. [3C ] How difficult is it to get DNA into a narrow tube? To an information theorist, the entropy associated with a constrained channel reveals how much information can be conveyed over it. In statistical physics, the same calculations are done for a different reason: to predict the thermodynamics of polymers, for example. As a toy example, consider a polymer of length N that can either sit in a constraining tube, of width L, or in the open where there are no constraints. In the open, the polymer adopts a state drawn at random from the set of one dimensional random walks, with, say, 3 possible directions per step. The entropy of this walk is log 3 per step, i.e., a total of N log 3. [The free energy of the polymer is defined to be −kT times this, where T is the temperature.] In the tube, the polymer’s onedimensional walk can go in 3 directions unless the wall is in the way, so the connection matrix is, for example (if L = 10),   1 1 0 0 0 0 0 0 0 0  1 1 1 0 0 0 0 0 0 0     0 1 1 1 0 0 0 0 0 0     0 0 1 1 1 0 0 0 0 0     0 0 0 1 1 1 0 0 0 0 .     .. .. ..   . . .    0 0 0 0 0 0 0 1 1 1  0 0 0 0 0 0 0 0 1 1 Now, what is the entropy of the polymer? What is the change in entropy associated with the polymer entering the tube? If possible, obtain an expression as a function of L. Use a computer to find the entropy of the walk for a particular value of L, e.g. 20, and plot the probability density of the polymer’s transverse location in the tube.

Notice the difference in capacity between two channels, one constrained and one unconstrained, is directly proportional to the force required to pull the DNA into the tube.

17.7 Solutions Solution to exercise 17.5 (p.250). A file transmitted by C 2 contains, on average, one-third 1s and two-thirds 0s. If f = 0.38, the fraction of 1s is f /(1 + f ) = (γ − 1.0)/(2γ − 1.0) = 0.2764. Solution to exercise 17.7 (p.254). A valid string for channel C can be obtained from a valid string for channel A by first inverting it [1 → 0; 0 → 1], then passing it through an accumulator. These operations are invertible, so any valid string for C can also be mapped onto a valid string for A. The only proviso here comes from the edge effects. If we assume that the first character transmitted over channel C is preceded by a string of zeroes, so that the first character is forced to be a 1 (figure 17.5c) then the two channels are exactly equivalent only if we assume that channel A’s first character must be a zero. Solution to exercise 17.8 (p.254). With N = 16 transmitted bits, the largest integer number of source bits that can be encoded is 10, so the maximum rate of a fixed length code with N = 16 is 0.625.

Figure 17.10. Model of DNA squashed in a narrow tube. The DNA will have a tendency to pop out of the tube, because, outside the tube, its random walk has greater entropy.

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258

17 — Communication over Constrained Noiseless Channels

Solution to exercise 17.10 (p.255).

Let the invariant distribution be

(R) P (s) = αe(L) s es ,

(17.22)

where α is a normalization constant. The entropy of S t given St−1 , assuming St−1 comes from the invariant distribution, is X H(St |St−1 ) = − P (s)P (s0 |s) log P (s0 |s) (17.23) s,s0

=

=−

X s,s0

α e(R) s



X

(L)

(R) αe(L) s es

s,s0

e s0 A s0 s (L)

λes

(L)

log

e s0 A s0 s (L)

λes

(L) i e s0 A s0 s h (L) . log es0 + log As0 s − log λ − log e(L) s λ

(17.24)

(17.25)

Now, As0 s is either 0 or 1, so the contributions from the terms proportional to As0 s log As0 s are all zero. So ! αX X (L) (L) (R) H(St |St−1 ) = log λ + − es0 log es0 + A s0 s e s λ 0 s s ! X X α (L) (L) (17.26) es0 As0 s e(R) s log es λ s 0 s

=

log λ −

=

log λ.

α X (R) (L) α X (L) (R) (L) λes es log e(L) (17.27) λes0 es0 log es0 + s λ 0 λ s s

(17.28)

Solution to exercise 17.11 (p.255). The principal eigenvalues of the connection matrices of the two channels are 1.839 and 1.928. The capacities (log λ) are 0.879 and 0.947 bits. Solution to exercise 17.12 (p.256). The channel is similar to the unconstrained binary channel; runs of length greater than L are rare if L is large, so we only expect weak differences from this channel; these differences will show up in contexts where the run length is close to L. The capacity of the channel is very close to one bit. A lower bound on the capacity is obtained by considering the simple variable-length code for this channel which replaces occurrences of the maximum runlength string 111. . .1 by 111. . .10, and otherwise leaves the source file unchanged. The average rate of this code is 1/(1 + 2 −L ) because the invariant distribution will hit the ‘add an extra zero’ state a fraction 2 −L of the time. We can reuse the solution for the variable-length channel in exercise 6.18 (p.125). The capacity is the value of β such that the equation Z(β) =

L+1 X

2−βl = 1

(17.29)

l=1

is satisfied. The L+1 terms in the sum correspond to the L+1 possible strings that can be emitted, 0, 10, 110, . . . , 11. . .10. The sum is exactly given by: Z(β) = 2

−β

L+1 2−β −1 . 2−β − 1

(17.30)

Here, as in Chapter 4, St denotes the ensemble whose random variable is the state st .

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17.7: Solutions "

259

# a(r N +1 − 1) Here we used . ar = r−1 n=0 We anticipate that β should be a little less than 1 in order for Z(β) to equal 1. Rearranging and solving approximately for β, using ln(1 + x) ' x, N X

n

Z(β) = 1

(17.31)

⇒ β ' 1 − 2−(L+2) / ln 2.

(17.32)

We evaluated the true capacities for L = 2 and L = 3 in an earlier exercise. The table compares the approximate capacity β with the true capacity for a selection of values of L. The element Q1|0 will be close to 1/2 (just a tiny bit larger), since in the unconstrained binary channel Q1|0 = 1/2. When a run of length L − 1 has occurred, we effectively have a choice of printing 10 or 0. Let the probability of selecting 10 be f . Let us estimate the entropy of the remaining N characters in the stream as a function of f , assuming the rest of the matrix Q to have been set to its optimal value. The entropy of the next N characters in the stream is the entropy of the first bit, H 2 (f ), plus the entropy of the remaining characters, which is roughly (N − 1) bits if we select 0 as the first bit and (N −2) bits if 1 is selected. More precisely, if C is the capacity of the channel (which is roughly 1), H(the next N chars) ' H2 (f ) + [(N − 1)(1 − f ) + (N − 2)f ] C

= H2 (f ) + N C − f C ' H2 (f ) + N − f. (17.33)

Differentiating and setting to zero to find the optimal f , we obtain: log2

1−f 1−f '1 ⇒ ' 2 ⇒ f ' 1/3. f f

(17.34)

The probability of emitting a 1 thus decreases from about 0.5 to about 1/3 as the number of emitted 1s increases. Here is the optimal matrix:  0 .3334 0 0 0 0 0 0 0 0  0 0 .4287 0 0 0 0 0 0 0    0 0 0 .4669 0 0 0 0 0 0    0 0 0 0 .4841 0 0 0 0 0    0 0 0 0 0 .4923 0 0 0 0  .  0 0 0 0 0 0 .4963 0 0 0    0 0 0 0 0 0 0 .4983 0 0    0 0 0 0 0 0 0 0 .4993 0    0 0 0 0 0 0 0 0 0 .4998  1 .6666 .5713 .5331 .5159 .5077 .5037 .5017 .5007 .5002 

Our rough theory works.

(17.35)

L

β

True capacity

2 3 4 5 6 9

0.910 0.955 0.977 0.9887 0.9944 0.9993

0.879 0.947 0.975 0.9881 0.9942 0.9993

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18 Crosswords and Codebreaking In this chapter we make a random walk through a few topics related to language modelling.

18.1 Crosswords The rules of crossword-making may be thought of as defining a constrained channel. The fact that many valid crosswords can be made demonstrates that this constrained channel has a capacity greater than zero. There are two archetypal crossword formats. In a ‘type A’ (or American) crossword, every row and column consists of a succession of words of length 2 or more separated by one or more spaces. In a ‘type B’ (or British) crossword, each row and column consists of a mixture of words and single characters, separated by one or more spaces, and every character lies in at least one word (horizontal or vertical). Whereas in a type A crossword every letter lies in a horizontal word and a vertical word, in a typical type B crossword only about half of the letters do so; the other half lie in one word only. Type A crosswords are harder to create than type B because of the constraint that no single characters are permitted. Type B crosswords are generally harder to solve because there are fewer constraints per character.

Why are crosswords possible? If a language has no redundancy, then any letters written on a grid form a valid crossword. In a language with high redundancy, on the other hand, it is hard to make crosswords (except perhaps a small number of trivial ones). The possibility of making crosswords in a language thus demonstrates a bound on the redundancy of that language. Crosswords are not normally written in genuine English. They are written in ‘word-English’, the language consisting of strings of words from a dictionary, separated by spaces. . Exercise 18.1.[2 ] Estimate the capacity of word-English, in bits per character. [Hint: think of word-English as defining a constrained channel (Chapter 17) and see exercise 6.18 (p.125).] The fact that many crosswords can be made leads to a lower bound on the entropy of word-English. For simplicity, we now model word-English by Wenglish, the language introduced in section 4.1 which consists of W words all of length L. The entropy of such a language, per character, including inter-word spaces, is: HW ≡

log 2 W . L+1

260

(18.1)

D U F F

A F A R

S T U D

T I T O

G I L D S

A D I E U

B A V P A L V A N C H J E

S T A T O S T H E R O V E M I S O O L S L T S S A H V E C O R U L R G L E I O T B D O A E S S R E

C I T E S U T T E R R O T

H M O E U P I M E R E A P P S T A I N L U C E A L E S C M A R M I R C A S T O T H E R T A E E P

E R B A K E O R I I A M E N T S M E L N T I N E S B E O E R H A P E U N I F E R S T O E T T N U T C R A T W O A A L B A T T L E E E E I S T L E S A U

S I S T E R K E N N Y

S I R E S

S A B R E

I R O N

L A T E

Y S E R

T E A R R A T A E R O H E L M

N G A R L L L E O A N

R O C K E T S

N

E X C U S E S

D U E S P H I E B R I

A L O H A

I E S L A T S O T N A O D E P S P K E R T A T T E

Y

R R N

Figure 18.1. Crosswords of types A (American) and B (British).

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18.1: Crosswords

261

We’ll find that the conclusions we come to depend on the value of H W and are not terribly sensitive to the value of L. Consider a large crossword of size S squares in area. Let the number of words be f w S and let the number of letter-occupied squares be f1 S. For typical crosswords of types A and B made of words of length L, the two fractions f w and f1 have roughly the values in table 18.2. We now estimate how many crosswords there are of size S using our simple model of Wenglish. We assume that Wenglish is created at random by generating W strings from a monogram (i.e., memoryless) source with entropy H 0 . If, for example, the source used all A = 26 characters with equal probability then H0 = log2 A = 4.7 bits. If instead we use Chapter 2’s distribution then the entropy is 4.2. The redundancy of Wenglish stems from two sources: it tends to use some letters more than others; and there are only W words in the dictionary. Let’s now count how many crosswords there are by imagining filling in the squares of a crossword at random using the same distribution that produced the Wenglish dictionary and evaluating the probability that this random scribbling produces valid words in all rows and columns. The total number of typical fillings-in of the f1 S squares in the crossword that can be made is |T | = 2f1 SH0 .

fw f1

A

B

2 L+1 L L+1

1 L+1 3 L 4L+1

Table 18.2. Factors fw and f1 by which the number of words and number of letter-squares respectively are smaller than the total number of squares.

(18.2)

The probability that one word of length L is validly filled-in is β=

W 2LH0

,

(18.3)

and the probability that the whole crossword, made of f w S words, is validly filled-in by a single typical in-filling is approximately β fw S .

(18.4)

So the log of the number of valid crosswords of size S is estimated to be log β fw S |T | = S [(f1 − fw L)H0 + fw log W ]

(18.5)

= S [(f1 − fw L)H0 + fw (L + 1)HW ] ,

(18.6)

which is an increasing function of S only if (f1 − fw L)H0 + fw (L + 1)HW > 0.

(18.7)

So arbitrarily many crosswords can be made only if there’s enough words in the Wenglish dictionary that HW >

(fw L − f1 ) H0 . fw (L + 1)

(18.8)

Plugging in the values of f1 and fw from table 18.2, we find the following. Crossword type Condition for crosswords

A HW >

1 L 2 L+1 H0

B HW >

1 L 4 L+1 H0

If we set H0 = 4.2 bits and assume there are W = 4000 words in a normal English-speaker’s dictionary, all with length L = 5, then we find that the condition for crosswords of type B is satisfied, but the condition for crosswords of type A is only just satisfied. This fits with my experience that crosswords of type A usually contain more obscure words.

This calculation underestimates the number of valid Wenglish crosswords by counting only crosswords filled with ‘typical’ strings. If the monogram distribution is non-uniform then the true count is dominated by ‘atypical’ fillings-in, in which crossword-friendly words appear more often.

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262

18 — Crosswords and Codebreaking

Further reading These observations about crosswords were first made by Shannon (1948); I learned about them from Wolf and Siegel (1998). The topic is closely related to the capacity of two-dimensional constrained channels. An example of a two-dimensional constrained channel is a two-dimensional bar-code, as seen on parcels. Exercise 18.2.[3 ] A two-dimensional channel is defined by the constraint that, of the eight neighbours of every interior pixel in an N × N rectangular grid, four must be black and four white. (The counts of black and white pixels around boundary pixels are not constrained.) A binary pattern satisfying this constraint is shown in figure 18.3. What is the capacity of this channel, in bits per pixel, for large N ?

Figure 18.3. A binary pattern in which every pixel is adjacent to four black and four white pixels.

18.2 Simple language models The Zipf–Mandelbrot distribution The crudest model for a language is the monogram model, which asserts that each successive word is drawn independently from a distribution over words. What is the nature of this distribution over words? Zipf’s law (Zipf, 1949) asserts that the probability of the rth most probable word in a language is approximately P (r) =

κ , rα

(18.9)

where the exponent α has a value close to 1, and κ is a constant. According to Zipf, a log–log plot of frequency versus word-rank should show a straight line with slope −α. Mandelbrot’s (1982) modification of Zipf’s law introduces a third parameter v, asserting that the probabilities are given by P (r) =

κ . (r + v)α

(18.10)

For some documents, such as Jane Austen’s Emma, the Zipf–Mandelbrot distribution fits well – figure 18.4. Other documents give distributions that are not so well fitted by a Zipf– Mandelbrot distribution. Figure 18.5 shows a plot of frequency versus rank for the LATEX source of this book. Qualitatively, the graph is similar to a straight line, but a curve is noticeable. To be fair, this source file is not written in pure English – it is a mix of English, maths symbols such as ‘x’, and LATEX commands. 0.1 to theand of I 0.01

is Harriet

0.001

information probability

0.0001

1e-05 1

10

100

1000

10000

Figure 18.4. Fit of the Zipf–Mandelbrot distribution (18.10) (curve) to the empirical frequencies of words in Jane Austen’s Emma (dots). The fitted parameters are κ = 0.56; v = 8.0; α = 1.26.

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18.2: Simple language models

263

0.1

Figure 18.5. Log–log plot of frequency versus rank for the words in the LATEX file of this book.

the of a is x 0.01 probability information 0.001 Shannon Bayes 0.0001

0.00001 1

10

100

1000

Figure 18.6. Zipf plots for four ‘languages’ randomly generated from Dirichlet processes with parameter α ranging from 1 to 1000. Also shown is the Zipf plot for this book.

alpha=1 0.1

alpha=10

0.01

alpha=100

0.001

alpha=1000

0.0001 book

0.00001

1

10

100

1000

10000

The Dirichlet process Assuming we are interested in monogram models for languages, what model should we use? One difficulty in modelling a language is the unboundedness of vocabulary. The greater the sample of language, the greater the number of words encountered. A generative model for a language should emulate this property. If asked ‘what is the next word in a newly-discovered work of Shakespeare?’ our probability distribution over words must surely include some non-zero probability for words that Shakespeare never used before. Our generative monogram model for language should also satisfy a consistency rule called exchangeability. If we imagine generating a new language from our generative model, producing an ever-growing corpus of text, all statistical properties of the text should be homogeneous: the probability of finding a particular word at a given location in the stream of text should be the same everywhere in the stream. The Dirichlet process model is a model for a stream of symbols (which we think of as ‘words’) that satisfies the exchangeability rule and that allows the vocabulary of symbols to grow without limit. The model has one parameter α. As the stream of symbols is produced, we identify each new symbol by a unique integer w. When we have seen a stream of length F symbols, we define the probability of the next symbol in terms of the counts {F w } of the symbols seen so far thus: the probability that the next symbol is a new symbol, never seen before, is α . (18.11) F +α The probability that the next symbol is symbol w is Fw . F +α

(18.12)

Figure 18.6 shows Zipf plots (i.e., plots of symbol frequency versus rank) for million-symbol ‘documents’ generated by Dirichlet process priors with values of α ranging from 1 to 1000. It is evident that a Dirichlet process is not an adequate model for observed distributions that roughly obey Zipf’s law.

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264

18 — Crosswords and Codebreaking 0.1

Figure 18.7. Zipf plots for the words of two ‘languages’ generated by creating successive characters from a Dirichlet process with α = 2, and declaring one character to be the space character. The two curves result from two different choices of the space character.

0.01

0.001

0.0001

0.00001 1

10

100

1000

10000

With a small tweak, however, Dirichlet processes can produce rather nice Zipf plots. Imagine generating a language composed of elementary symbols using a Dirichlet process with a rather small value of the parameter α, so that the number of reasonably frequent symbols is about 27. If we then declare one of those symbols (now called ‘characters’ rather than words) to be a space character, then we can identify the strings between the space characters as ‘words’. If we generate a language in this way then the frequencies of words often come out as very nice Zipf plots, as shown in figure 18.7. Which character is selected as the space character determines the slope of the Zipf plot – a less probable space character gives rise to a richer language with a shallower slope.

18.3 Units of information content The information content of an outcome, x, whose probability is P (x), is defined to be 1 h(x) = log . (18.13) P (x) The entropy of an ensemble is an average information content, H(X) =

X x

P (x) log

1 . P (x)

(18.14)

When we compare hypotheses with each other in the light of data, it is often convenient to compare the log of the probability of the data under the alternative hypotheses, ‘log evidence for Hi ’ = log P (D | Hi ),

(18.15)

or, in the case where just two hypotheses are being compared, we evaluate the ‘log odds’, P (D | H1 ) log , (18.16) P (D | H2 )

which has also been called the ‘weight of evidence in favour of H 1 ’. The log evidence for a hypothesis, log P (D | H i ) is the negative of the information content of the data D: if the data have large information content, given a hypothesis, then they are surprising to that hypothesis; if some other hypothesis is not so surprised by the data, then that hypothesis becomes more probable. ‘Information content’, ‘surprise value’, and log likelihood or log evidence are the same thing. All these quantities are logarithms of probabilities, or weighted sums of logarithms of probabilities, so they can all be measured in the same units. The units depend on the choice of the base of the logarithm. The names that have been given to these units are shown in table 18.8.

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18.4: A taste of Banburismus

265

Unit

Expression that has those units

bit nat ban deciban (db)

log2 p log e p log 10 p 10 log 10 p

The bit is the unit that we use most in this book. Because the word ‘bit’ has other meanings, a backup name for this unit is the shannon. A byte is 8 bits. A megabyte is 220 ' 106 bytes. If one works in natural logarithms, information contents and weights of evidence are measured in nats. The most interesting units are the ban and the deciban.

The history of the ban Let me tell you why a factor of ten in probability is called a ban. When Alan Turing and the other codebreakers at Bletchley Park were breaking each new day’s Enigma code, their task was a huge inference problem: to infer, given the day’s cyphertext, which three wheels were in the Enigma machines that day; what their starting positions were; what further letter substitutions were in use on the steckerboard; and, not least, what the original German messages were. These inferences were conducted using Bayesian methods (of course!), and the chosen units were decibans or half-decibans, the deciban being judged the smallest weight of evidence discernible to a human. The evidence in favour of particular hypotheses was tallied using sheets of paper that were specially printed in Banbury, a town about 30 miles from Bletchley. The inference task was known as Banburismus, and the units in which Banburismus was played were called bans, after that town.

18.4 A taste of Banburismus The details of the code-breaking methods of Bletchley Park were kept secret for a long time, but some aspects of Banburismus can be pieced together. I hope the following description of a small part of Banburismus is not too inaccurate.1 How much information was needed? The number of possible settings of the Enigma machine was about 8 × 1012 . To deduce the state of the machine, ‘it was therefore necessary to find about 129 decibans from somewhere’, as Good puts it. Banburismus was aimed not at deducing the entire state of the machine, but only at figuring out which wheels were in use; the logic-based bombes, fed with guesses of the plaintext (cribs), were then used to crack what the settings of the wheels were. The Enigma machine, once its wheels and plugs were put in place, implemented a continually-changing permutation cypher that wandered deterministically through a state space of 26 3 permutations. Because an enormous number of messages were sent each day, there was a good chance that whatever state one machine was in when sending one character of a message, there would be another machine in the same state while sending a particular character in another message. Because the evolution of the machine’s state was deterministic, the two machines would remain in the same state as each other 1 I’ve been most helped by descriptions given by Tony Sale (http://www. codesandciphers.org.uk/lectures/) and by Jack Good (1979), who worked with Turing at Bletchley.

Table 18.8. Units of measurement of information content.

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18 — Crosswords and Codebreaking

for the rest of the transmission. The resulting correlations between the outputs of such pairs of machines provided a dribble of information-content from which Turing and his co-workers extracted their daily 129 decibans.

How to detect that two messages came from machines with a common state sequence The hypotheses are the null hypothesis, H 0 , which states that the machines are in different states, and that the two plain messages are unrelated; and the ‘match’ hypothesis, H1 , which says that the machines are in the same state, and that the two plain messages are unrelated. No attempt is being made here to infer what the state of either machine is. The data provided are the two cyphertexts x and y; let’s assume they both have length T and that the alphabet size is A (26 in Enigma). What is the probability of the data, given the two hypotheses? First, the null hypothesis. This hypothesis asserts that the two cyphertexts are given by x = x1 x2 x3 . . . = c1 (u1 )c2 (u2 )c3 (u3 ) . . . (18.17) and y = y1 y2 y3 . . . = c01 (v1 )c02 (v2 )c03 (v3 ) . . . ,

(18.18)

where the codes ct and c0t are two unrelated time-varying permutations of the alphabet, and u1 u2 u3 . . . and v1 v2 v3 . . . are the plaintext messages. An exact computation of the probability of the data (x, y) would depend on a language model of the plain text, and a model of the Enigma machine’s guts, but if we assume that each Enigma machine is an ideal random time-varying permutation, then the probability distribution of the two cyphertexts is uniform. All cyphertexts are equally likely. P (x, y | H0 ) =

 2T 1 for all x, y of length T . A

(18.19)

What about H1 ? This hypothesis asserts that a single time-varying permutation ct underlies both x = x1 x2 x3 . . . = c1 (u1 )c2 (u2 )c3 (u3 ) . . .

(18.20)

y = y1 y2 y3 . . . = c1 (v1 )c2 (v2 )c3 (v3 ) . . . .

(18.21)

and What is the probability of the data (x, y)? We have to make some assumptions about the plaintext language. If it were the case that the plaintext language was completely random, then the probability of u 1 u2 u3 . . . and v1 v2 v3 . . . would be uniform, and so would that of x and y, so the probability P (x, y | H 1 ) would be equal to P (x, y | H0 ), and the two hypotheses H0 and H1 would be indistinguishable. We make progress by assuming that the plaintext is not completely random. Both plaintexts are written in a language, and that language has redundancies. Assume for example that particular plaintext letters are used more often than others. So, even though the two plaintext messages are unrelated, they are slightly more likely to use the same letters as each other; if H 1 is true, two synchronized letters from the two cyphertexts are slightly more likely to be identical. Similarly, if a language uses particular bigrams and trigrams frequently, then the two plaintext messages will occasionally contain the same bigrams and trigrams at the same time as each other, giving rise, if H 1 is true,

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18.4: A taste of Banburismus u v matches:

267

LITTLE-JACK-HORNER-SAT-IN-THE-CORNER-EATING-A-CHRISTMAS-PIE--HE-PUT-IN-H RIDE-A-COCK-HORSE-TO-BANBURY-CROSS-TO-SEE-A-FINE-LADY-UPON-A-WHITE-HORSE .*....*..******.*..............*...........*................*...........

to a little burst of 2 or 3 identical letters. Table 18.9 shows such a coincidence in two plaintext messages that are unrelated, except that they are both written in English. The codebreakers hunted among pairs of messages for pairs that were suspiciously similar to each other, counting up the numbers of matching monograms, bigrams, trigrams, etc. This method was first used by the Polish codebreaker Rejewski. Let’s look at the simple case of a monogram language model and estimate how long a message is needed to be able to decide whether two machines are in the same state. I’ll assume the source language is monogram-English, the language in which successive letters are drawn i.i.d. from the probability distribution {pi } of figure 2.1. The probability of x and y is nonuniform: consider two single characters, xt = ct (ut ) and yt = ct (vt ); the probability that they are identical is X X p2i ≡ m. (18.22) P (ut )P (vt ) [ut = vt ] = 

ut ,vt

i

We give this quantity the name m, for ‘match probability’; for both English and German, m is about 2/26 rather than 1/26 (the value that would hold for a completely random language). Assuming that c t is an ideal random permutation, the probability of xt and yt is, by symmetry, ( m if xt = yt A P (xt , yt | H1 ) = (18.23) (1−m) for xt 6= yt . A(A−1) Given a pair of cyphertexts x and y of length T that match in M places and do not match in N places, the log evidence in favour of H 1 is then P (x, y | H1 ) log P (x, y | H0 )

(1−m)

m/A A(A−1) = M log + N log 1/A2 1/A2 (1 − m)A . = M log mA + N log A−1

(18.24) (18.25)

Every match contributes log mA in favour of H 1 ; every non-match contributes A−1 in favour of H0 . log (1−m)A Match probability for monogram-English Coincidental match probability log-evidence for H1 per match log-evidence for H1 per non-match

m 1/A 10 log 10 mA 10 log 10 (1−m)A (A−1)

0.076 0.037 3.1 db −0.18 db

If there were M = 4 matches and N = 47 non-matches in a pair of length T = 51, for example, the weight of evidence in favour of H 1 would be +4 decibans, or a likelihood ratio of 2.5 to 1 in favour. The expected weight of evidence from a line of text of length T = 20 characters is the expectation of (18.25), which depends on whether H 1 or H0 is true. If H1 is true then matches are expected to turn up at rate m, and the expected weight of evidence is 1.4 decibans per 20 characters. If H 0 is true

Table 18.9. Two aligned pieces of English plaintext, u and v, with matches marked by *. Notice that there are twelve matches, including a run of six, whereas the expected number of matches in two completely random strings of length T = 74 would be about 3. The two corresponding cyphertexts from two machines in identical states would also have twelve matches.

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268

18 — Crosswords and Codebreaking

then spurious matches are expected to turn up at rate 1/A, and the expected weight of evidence is −1.1 decibans per 20 characters. Typically, roughly 400 characters need to be inspected in order to have a weight of evidence greater than a hundred to one (20 decibans) in favour of one hypothesis or the other. So, two English plaintexts have more matches than two random strings. Furthermore, because consecutive characters in English are not independent, the bigram and trigram statistics of English are nonuniform and the matches tend to occur in bursts of consecutive matches. [The same observations also apply to German.] Using better language models, the evidence contributed by runs of matches was more accurately computed. Such a scoring system was worked out by Turing and refined by Good. Positive results were passed on to automated and human-powered codebreakers. According to Good, the longest false-positive that arose in this work was a string of 8 consecutive matches between two machines that were actually in unrelated states.

Further reading For further reading about Turing and Bletchley Park, see Hodges (1983) and Good (1979). For an in-depth read about cryptography, Schneier’s (1996) book is highly recommended. It is readable, clear, and entertaining.

18.5 Exercises . Exercise 18.3.[2 ] Another weakness in the design of the Enigma machine, which was intended to emulate a perfectly random time-varying permutation, is that it never mapped a letter to itself. When you press Q, what comes out is always a different letter from Q. How much information per character is leaked by this design flaw? How long a crib would be needed to be confident that the crib is correctly aligned with the cyphertext? And how long a crib would be needed to be able confidently to identify the correct key? [A crib is a guess for what the plaintext was. Imagine that the Brits know that a very important German is travelling from Berlin to Aachen, and they intercept Enigma-encoded messages sent to Aachen. It is a good bet that one or more of the original plaintext messages contains the string OBERSTURMBANNFUEHRERXGRAFXHEINRICHXVONXWEIZSAECKER, the name of the important chap. A crib could be used in a brute-force approach to find the correct Enigma key (feed the received messages through all possible Engima machines and see if any of the putative decoded texts match the above plaintext). This question centres on the idea that the crib can also be used in a much less expensive manner: slide the plaintext crib along all the encoded messages until a perfect mismatch of the crib and the encoded message is found; if correct, this alignment then tells you a lot about the key.]

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19 Why have Sex? Information Acquisition and Evolution Evolution has been happening on earth for about the last 10 9 years. Undeniably, information has been acquired during this process. Thanks to the tireless work of the Blind Watchmaker, some cells now carry within them all the information required to be outstanding spiders; other cells carry all the information required to make excellent octopuses. Where did this information come from? The entire blueprint of all organisms on the planet has emerged in a teaching process in which the teacher is natural selection: fitter individuals have more progeny, the fitness being defined by the local environment (including the other organisms). The teaching signal is only a few bits per individual: an individual simply has a smaller or larger number of grandchildren, depending on the individual’s fitness. ‘Fitness’ is a broad term that could cover • the ability of an antelope to run faster than other antelopes and hence avoid being eaten by a lion; • the ability of a lion to be well-enough camouflaged and run fast enough to catch one antelope per day; • the ability of a peacock to attract a peahen to mate with it; • the ability of a peahen to rear many young simultaneously. The fitness of an organism is largely determined by its DNA – both the coding regions, or genes, and the non-coding regions (which play an important role in regulating the transcription of genes). We’ll think of fitness as a function of the DNA sequence and the environment. How does the DNA determine fitness, and how does information get from natural selection into the genome? Well, if the gene that codes for one of an antelope’s proteins is defective, that antelope might get eaten by a lion early in life and have only two grandchildren rather than forty. The information content of natural selection is fully contained in a specification of which offspring survived to have children – an information content of at most one bit per offspring. The teaching signal does not communicate to the ecosystem any description of the imperfections in the organism that caused it to have fewer children. The bits of the teaching signal are highly redundant, because, throughout a species, unfit individuals who are similar to each other will be failing to have offspring for similar reasons. So, how many bits per generation are acquired by the species as a whole by natural selection? How many bits has natural selection succeeded in conveying to the human branch of the tree of life, since the divergence between 269

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270

19 — Why have Sex? Information Acquisition and Evolution

Australopithecines and apes 4 000 000 years ago? Assuming a generation time of 10 years for reproduction, there have been about 400 000 generations of human precursors since the divergence from apes. Assuming a population of 109 individuals, each receiving a couple of bits of information from natural selection, the total number of bits of information responsible for modifying the genomes of 4 million B.C. into today’s human genome is about 8 × 10 14 bits. However, as we noted, natural selection is not smart at collating the information that it dishes out to the population, and there is a great deal of redundancy in that information. If the population size were twice as great, would it evolve twice as fast? No, because natural selection will simply be correcting the same defects twice as often. John Maynard Smith has suggested that the rate of information acquisition by a species is independent of the population size, and is of order 1 bit per generation. This figure would allow for only 400 000 bits of difference between apes and humans, a number that is much smaller than the total size of the human genome – 6 × 109 bits. [One human genome contains about 3 × 10 9 nucleotides.] It is certainly the case that the genomic overlap between apes and humans is huge, but is the difference that small? In this chapter, we’ll develop a crude model of the process of information acquisition through evolution, based on the assumption that a gene with two defects is typically likely to be more defective than a gene with one defect, and an organism with two defective genes is likely to be less fit than an organism with one defective gene. Undeniably, this is a crude model, since real biological systems are baroque constructions with complex interactions. Nevertheless, we persist with a simple model because it readily yields striking results. What we find from this simple model is that 1. John Maynard Smith’s figure of 1 bit per generation is correct for an asexually-reproducing population; 2. in contrast, if the species reproduces sexually, the rate of information √ acquisition can be as large as G bits per generation, where G is the size of the genome. We’ll also find interesting results concerning the maximum mutation rate that a species can withstand.

19.1 The model We study a simple model of a reproducing population of N individuals with a genome of size G bits: variation is produced by mutation or by recombination (i.e., sex) and truncation selection selects the N fittest children at each generation to be the parents of the next. We find striking differences between populations that have recombination and populations that do not. The genotype of each individual is a vector x of G bits, each having a good state xg = 1 and a bad state xg = 0. The fitness F (x) of an individual is simply the sum of her bits: G X F (x) = xg . (19.1) g=1

The bits in the genome could be considered to correspond either to genes that have good alleles (xg = 1) and bad alleles (xg = 0), or to the nucleotides of a genome. We will concentrate on the latter interpretation. The essential property of fitness that we are assuming is that it is locally a roughly linear function of the genome, that is, that there are many possible changes one

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19.2: Rate of increase of fitness

271

could make to the genome, each of which has a small effect on fitness, and that these effects combine approximately linearly. We define the normalized fitness f (x) ≡ F (x)/G. We consider evolution by natural selection under two models of variation. Variation by mutation. The model assumes discrete generations. At each generation, t, every individual produces two children. The children’s genotypes differ from the parent’s by random mutations. Natural selection selects the fittest N progeny in the child population to reproduce, and a new generation starts. [The selection of the fittest N individuals at each generation is known as truncation selection.] The simplest model of mutations is that the child’s bits {x g } are independent. Each bit has a small probability of being flipped, which, thinking of the bits as corresponding roughly to nucleotides, is taken to be a constant m, independent of xg . [If alternatively we thought of the bits as corresponding to genes, then we would model the probability of the discovery of a good gene, P (xg = 0 → xg = 1), as being a smaller number than the probability of a deleterious mutation in a good gene, P (xg = 1 → xg = 0).] Variation by recombination (or crossover, or sex). Our organisms are haploid, not diploid. They enjoy sex by recombination. The N individuals in the population are married into M = N/2 couples, at random, and each couple has C children – with C = 4 children being our standard assumption, so as to have the population double and halve every generation, as before. The C children’s genotypes are independent given the parents’. Each child obtains its genotype z by random crossover of its parents’ genotypes, x and y. The simplest model of recombination has no linkage, so that:  xg with probability 1/2 (19.2) zg = yg with probability 1/2. Once the M C progeny have been born, the parents pass away, the fittest N progeny are selected by natural selection, and a new generation starts. We now study these two models of variation in detail.

19.2 Rate of increase of fitness Theory of mutations We assume that the genotype of an individual with normalized fitness f = F/G is subjected to mutations that flip bits with probability m. We first show that if the average normalized fitness f of the population is greater than 1/2, then the optimal mutation rate is small, and the rate of acquisition of information is at most of order one bit per generation. Since it is easy to achieve a normalized fitness of f = 1/2 by simple mutation, we’ll assume f > 1/2 and work in terms of the excess normalized fitness δf ≡ f − 1/2. If an individual with excess normalized fitness δf has a child and the mutation rate m is small, the probability distribution of the excess normalized fitness of the child has mean δf child = (1 − 2m)δf

(19.3)

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272

19 — Why have Sex? Information Acquisition and Evolution

and variance

m(1 − m) m ' . G G

(19.4)

If the population of parents has mean δf (t) and variance σ 2 (t) ≡ βm/G, then the child population, before selection, will have mean (1 − 2m)δf (t) and variance (1+β)m/G. Natural selection chooses the upper half of this distribution, so the mean fitness and variance of fitness at the next generation are given by r p m , (19.5) δf (t+1) = (1 − 2m)δf (t) + α (1 + β) G

m , (19.6) G where α is the mean deviation from the mean, measured in standard deviations, and γ is the factor by which the child distribution’s variance is reduced by selection. The numbers α and γ are of order 1. For the case of a Gaussian p distribution, α = 2/π ' 0.8 and γ = (1 − 2/π) ' 0.36. If we assume that the variance is in dynamic equilibrium, i.e., σ 2 (t+1) ' σ 2 (t), then σ 2 (t+1) = γ(1 + β)

γ(1 + β) = β, so (1 + β) =

1 , 1−γ

(19.7)

p and the factor α (1 + β) in equation (19.5) is equal to 1, if we take the results for the Gaussian distribution, an approximation that becomes poorest when the discreteness of fitness becomes important, i.e., for small m. The rate of increase of normalized fitness is thus: r df m ' −2m δf + , (19.8) dt G which, assuming G(δf )2  1, is maximized for mopt = at which point,



df dt



opt

1 , 16G(δf )2

=

1 . 8G(δf )

(19.9)

(19.10)

So the rate of increase of fitness F = f G is at most dF 1 = per generation. dt 8(δf )

(19.11)

For a population with low fitness (δf < 0.125), the√rate of increase of fitness may exceed 1 unit per √ generation. Indeed, if δf 1/ G, the rate √ of increase, if 1 / m = 2, is of order G; this initial spurt can last only of order G generations. For δf > 0.125, the rate of increase of fitness is smaller than one per generation. As the fitness approaches G, the optimal mutation rate tends to m = 1/(4G), so that an average of 1/4 bits are flipped per genotype, and the rate of increase of fitness is also equal to 1/4; information is gained at a rate of about 0.5 bits per generation. It takes about 2G generations for the genotypes of all individuals in the population to attain perfection. For fixed m, the fitness is given by 1 (1 − c e−2mt ), δf (t) = √ 2 mG

(19.12)

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19.2: Rate of increase of fitness

273 No sex

Sex

Histogram of parents’ fitness

Histogram of children’s fitness

Selected children’s fitness

subject to the constraint δf (t) ≤ 1/2, where c is a constant of integration, equal to 1 if f (0) = 1/2. If the mean number of bits flipped per genotype, mG, exceeds √ 1, then the fitness F approaches an equilibrium value F eqm = (1/2 + 1/(2 mG))G. This theory is somewhat inaccurate in that the true probability distribution of fitness is non-Gaussian, asymmetrical, and quantized to integer values. All the same, the predictions of the theory are not grossly at variance with the results of simulations described below.

Theory of sex The analysis of the sexual population becomes tractable with two approximations: first, we assume that the gene-pool mixes sufficiently rapidly that correlations between genes can be neglected; second, we assume homogeneity, i.e., that the fraction fg of bits g that are in the good state is the same, f (t), for all g. Given these assumptions, if two parents of fitness F = f G mate, the probability distribution of their children’s fitness has mean equal to the parents’ fitness, F ; the variation produced by sex does not reduce the average fitness. p The standard deviation of the fitness of the children scales as Gf (1 − f ). Since, after selection, the increase in fitness is proportional to this standard deviation, the fitness √ increase per generation scales as the square root of the size of the genome, G. As shown in box 19.2, the mean fitness F¯ = f G evolves in accordance with the differential equation: p dF¯ ' η f (t)(1 − f (t))G, dt

(19.13)

p where η ≡ 2/(π + 2). The solution of this equation is      √ √ η 1 f (t) = 1 + sin √ (t + c) , for t + c ∈ − π2 G/η, π2 G/η , (19.14) 2 G

where c is a constant of integration, c = sin −1 (2f (0) − 1). So this idealized system √ reaches a state of eugenic perfection (f = 1) within a finite time: (π/η) G generations.

Simulations Figure 19.3a shows the fitness of a sexual population of N = 1000 individuals with a genome size of G = 1000 starting from a random initial state with normalized fitness 0.5. It also shows the theoretical curve f (t)G from equation (19.14), which fits remarkably well. In contrast, figures 19.3(b) and (c) show the evolving fitness when variation is produced by mutation at rates m = 0.25/G and m = 6/G respectively. Note the difference in the horizontal scales from panel (a).

Figure 19.1. Why sex is better than sex-free reproduction. If mutations are used to create variation among children, then it is unavoidable that the average fitness of the children is lower than the parents’ fitness; the greater the variation, the greater the average deficit. Selection bumps up the mean fitness again. In contrast, recombination produces variation without a decrease in average fitness. The typical √ amount of variation scales as G, where G is the genome size, so after selection, √ the average fitness rises by O( G).

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274

19 — Why have Sex? Information Acquisition and Evolution

How does f (t+1) depend on f (t)? Let’s first assume the two parents of a child both have exactly f (t)G good bits, and, by our homogeneity assumption, that those bits are independent random subsets of the G bits. The number of bits that are good in both parents is roughly f (t)2 G, and the number that are good in one parent only is roughly 2f (t)(1−f (t))G, so the fitness of the child will be f (t)2 G plus the sum of 2f (t)(1−f (t))G fair coin flips, which has a binomial distribution of mean f (t)(1 − f (t))G and variance 1 f (t)(1 − f (t))G. The fitness of a child is thus roughly distributed as 2 1 Fchild ∼ Normal mean = f (t)G, variance = f (t)(1 − f (t))G . 2 

The important property of this distribution, contrasted with the distribution under mutation, is that the mean fitness is equal to the parents’ fitness; the variation produced by sex does not reduce the average fitness. If we include the parental population’s variance, which we will write as σ 2 (t) = β(t) 21 f (t)(1 − f (t))G, the children’s fitnesses are distributed as Fchild ∼ Normal mean = f (t)G, variance =

1+

β 2 

1 f (t)(1 − f (t))G . 2 

Natural selection selects the children on the upper side of this distribution. The mean increase in fitness will be √ F¯ (t+1) − F¯ (t) = [α(1 + β/2)1/2 / 2] f (t)(1 − f (t))G, 

and the variance of the surviving children will be 1 σ 2 (t + 1) = γ(1 + β/2) f (t)(1 − f (t))G, 2 where α = 2/π and γ = (1 − 2/π). If there is dynamic equilibrium [σ 2 (t + 1) = σ 2 (t)] then the factor in (19.2) is 

√ α(1 + β/2)1/2 / 2 = 

2 ' 0.62. (π + 2)

Defining this constant to be η ≡ 2/(π + 2), we conclude that, under sex and natural selection, the mean fitness of the population increases at a rate proportional to the square root of the size of the genome, 

dF¯ 'η dt 

f (t)(1 − f (t))G bits per generation.

Box 19.2. Details of the theory of sex.

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19.3: The maximal tolerable mutation rate

275 Figure 19.3. Fitness as a function of time. The genome size is G = 1000. The dots show the fitness of six randomly selected individuals from the birth population at each generation. The initial population of N = 1000 had randomly generated genomes with f (0) = 0.5 (exactly). (a) Variation produced by sex alone. Line shows theoretical curve (19.14) for infinite homogeneous population. (b,c) Variation produced by mutation, with and without sex, when the mutation rate is mG = 0.25 (b) or 6 (c) bits per genome. The dashed line shows the curve (19.12).

1000 900 800 700 600 500

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Exercise 19.1.[3, p.280] Dependence on population size. How do the results for a sexual population depend on the population size? We anticipate that there is a minimum population size above which the theory of sex is accurate. How is that minimum population size related to G? Exercise 19.2.[3 ] Dependence on crossover mechanism. In the simple model of sex, each bit is taken at random from one of the two parents, that is, we allow crossovers to occur with probability 50% between any two adjacent nucleotides. How is the model affected (a) if the crossover probability is smaller? (b) if crossovers occur exclusively at hot-spots located every d bits along the genome?

19.3 The maximal tolerable mutation rate What if we combine the two models of variation? What is the maximum mutation rate that can be tolerated by a species that has sex? The rate of increase of fitness is given by r √ df m + f (1 − f )/2 ' −2m δf + η 2 , (19.15) dt G

Figure 19.4. Maximal tolerable mutation rate, shown as number of errors per genome (mG), versus normalized fitness f = F/G. Left panel: genome size G = 1000; right: G = 100 000. Independent of genome size, a parthenogenetic species (no sex) can tolerate only of order 1 error per genome per generation; a species that uses recombination (sex) can tolerate far greater mutation rates.

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19 — Why have Sex? Information Acquisition and Evolution

which is positive if the mutation rate satisfies r f (1 − f ) . (19.16) m<η G Let us compare this rate with the result in the absence of sex, which, from equation (19.8), is that the maximum tolerable mutation rate is 1 1 m< . (19.17) G (2 δf )2 √ The tolerable mutation rate with sex is of order G times greater than that without sex! A parthenogenetic (non-sexual) species could try to wriggle out of this bound on its mutation rate by increasing its litter sizes. But if mutation flips on average mG bits, the probability that no bits are flipped in one genome is roughly e−mG , so a mother needs to have roughly emG offspring in order to have a good chance of having one child with the same fitness as her. The litter size of a non-sexual species thus has to be exponential in mG (if mG is bigger than 1), if the species is to persist. So the maximum tolerable mutation rate is pinned close √ to 1/G, for a nonsexual species, whereas it is a larger number of order 1/ G, for a species with recombination. Turning these results around, we can predict the largest possible genome size for a given fixed mutation rate, m. For a parthenogenetic species, the largest genome size is of order 1/m, and for a sexual species, 1/m 2 . Taking the figure m = 10−8 as the mutation rate per nucleotide per generation (EyreWalker and Keightley, 1999), and allowing for a maximum brood size of 20 000 (that is, mG ' 10), we predict that all species with more than G = 10 9 coding nucleotides make at least occasional use of recombination. If the brood size is 12, then this number falls to G = 2.5 × 10 8 .

19.4 Fitness increase and information acquisition For this simple model it is possible to relate increasing fitness to information acquisition. If the bits are set at random, the fitness is roughly F = G/2. If evolution leads to a population in which all individuals have the maximum fitness F = G, then G bits of information have been acquired by the species, namely for each bit xg , the species has figured out which of the two states is the better. We define the information acquired at an intermediate fitness to be the amount of selection (measured in bits) required to select the perfect state from the gene pool. Let a fraction fg of the population have xg = 1. Because log 2 (1/f ) is the information required to find a black ball in an urn containing black and white balls in the ratio f : 1−f , we define the information acquired to be X fg I= bits. (19.18) log2 1/2 g If all the fractions fg are equal to F/G, then I = G log 2

2F , G

(19.19)

which is well approximated by I˜ ≡ 2(F − G/2).

(19.20)

The rate of information acquisition is thus roughly two times the rate of increase of fitness in the population.

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19.5: Discussion

19.5 Discussion These results quantify the well known argument for why species reproduce by sex with recombination, namely that recombination allows useful mutations to spread more rapidly through the species and allows deleterious mutations to be more rapidly cleared from the population (Maynard Smith, 1978; Felsenstein, 1985; Maynard Smith, 1988; Maynard Smith and Sz´athmary, 1995). A population that reproduces by recombination can acquire informa√ tion from natural selection at a rate of order G times faster than a partheno√ genetic population, and it can tolerate a mutation rate that is of order G times greater. For genomes of size G ' 10 8 coding nucleotides, this factor of √ G is substantial. This enormous advantage conferred by sex has been noted before by Kondrashov (1988), but this meme, which Kondrashov calls ‘the deterministic mutation hypothesis’, does not seem to have diffused throughout the evolutionary research community, as there are still numerous papers in which the prevalence of sex is viewed as a mystery to be explained by elaborate mechanisms.

‘The cost of males’ – stability of a gene for sex or parthenogenesis Why do people declare sex to be a mystery? The main motivation for being mystified is an idea called the ‘cost of males’. Sexual reproduction is disadvantageous compared with asexual reproduction, it’s argued, because of every two offspring produced by sex, one (on average) is a useless male, incapable of child-bearing, and only one is a productive female. In the same time, a parthenogenetic mother could give birth to two female clones. To put it another way, the big advantage of parthenogenesis, from the point of view of the individual, is that one is able to pass on 100% of one’s genome to one’s children, instead of only 50%. Thus if there were two versions of a species, one reproducing with and one without sex, the single mothers would be expected to outstrip their sexual cousins. The simple model presented thus far did not include either genders or the ability to convert from sexual reproduction to asexual, but we can easily modify the model. We modify the model so that one of the G bits in the genome determines whether an individual prefers to reproduce parthenogenetically (x = 1) or sexually (x = 0). The results depend on the number of children had by a single parthenogenetic mother, Kp and the number of children born by a sexual couple, Ks . Both (Kp = 2, Ks = 4) and (Kp = 4, Ks = 4) are reasonable models. The former (Kp = 2, Ks = 4) would seem most appropriate in the case of unicellular organisms, where the cytoplasm of both parents goes into the children. The latter (Kp = 4, Ks = 4) is appropriate if the children are solely nurtured by one of the parents, so single mothers have just as many offspring as a sexual pair. I concentrate on the latter model, since it gives the greatest advantage to the parthenogens, who are supposedly expected to outbreed the sexual community. Because parthenogens have four children per generation, the maximum tolerable mutation rate for them is twice the expression (19.17) derived before for Kp = 2. If the fitness is large, the maximum tolerable rate is mG ' 2. Initially the genomes are set randomly with F = G/2, with half of the population having the gene for parthenogenesis. Figure 19.5 shows the outcome. During the ‘learning’ phase of evolution, in which the fitness is increasing rapidly, pockets of parthenogens appear briefly, but then disappear within a couple of generations as their sexual cousins overtake them in fitness and

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19 — Why have Sex? Information Acquisition and Evolution

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leave them behind. Once the population reaches its top fitness, however, the parthenogens can take over, if the mutation rate is sufficiently low (mG = 1). In the presence of a higher mutation rate (mG = 4), however, the parthenogens never √ take over. The breadth of the sexual population’s fitness is of order G, so a mutant parthenogenetic colony arising √ √ with slightly above-average fitness will last for about G/(mG) = 1/(m G) generations before its fitness falls below that of its sexual cousins. As long as the population size is sufficiently large for some sexual individuals to survive for this time, sex will not die out. In a sufficiently unstable environment, where the fitness function is continually changing, the parthenogens will always lag behind the sexual community. These results are consistent with the argument of Haldane and Hamilton (2002) that sex is helpful in an arms race with parasites. The parasites define an effective fitness function which changes with time, and a sexual population will always ascend the current fitness function more rapidly.

Additive fitness function Of course, our results depend on the fitness function that we assume, and on our model of selection. Is it reasonable to model fitness, to first order, as a sum of independent terms? Maynard Smith (1968) argues that it is: the more good genes you have, the higher you come in the pecking order, for example. The directional selection model has been used extensively in theoretical population genetic studies (Bulmer, 1985). We might expect real fitness functions to involve interactions, in which case crossover might reduce the average fitness. However, since recombination gives the biggest advantage to species whose fitness functions are additive, we might predict that evolution will have favoured species that used a representation of the genome that corresponds to a fitness function that has only weak interactions. And even if there are interactions, it seems plausible that the fitness would still involve a sum of such interacting terms, with the number of terms being some fraction of the genome size G.

Figure 19.5. Results when there is a gene for parthenogenesis, and no interbreeding, and single mothers produce as many children as sexual couples. G = 1000, N = 1000. (a) mG = 4; (b) mG = 1. Vertical axes show the fitnesses of the two sub-populations, and the percentage of the population that is parthenogenetic.

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19.6: Further exercises Exercise 19.3.[3C ] Investigate how fast sexual and asexual species evolve if they have a fitness function with interactions. For example, let the fitness be a sum of exclusive-ors of pairs of bits; compare the evolving fitnesses with those of the sexual and asexual species with a simple additive fitness function. Furthermore, if the fitness function were a highly nonlinear function of the genotype, it could be made more smooth and locally linear by the Baldwin effect. The Baldwin effect (Baldwin, 1896; Hinton and Nowlan, 1987) has been widely studied as a mechanism whereby learning guides evolution, and it could also act at the level of transcription and translation. Consider the evolution of a peptide sequence for a new purpose. Assume the effectiveness of the peptide is a highly nonlinear function of the sequence, perhaps having a small island of good sequences surrounded by an ocean of equally bad sequences. In an organism whose transcription and translation machinery is flawless, the fitness will be an equally nonlinear function of the DNA sequence, and evolution will wander around the ocean making progress towards the island only by a random walk. In contrast, an organism having the same DNA sequence, but whose DNA-to-RNA transcription or RNA-to-protein translation is ‘faulty’, will occasionally, by mistranslation or mistranscription, accidentally produce a working enzyme; and it will do so with greater probability if its DNA sequence is close to a good sequence. One cell might produce 1000 proteins from the one mRNA sequence, of which 999 have no enzymatic effect, and one does. The one working catalyst will be enough for that cell to have an increased fitness relative to rivals whose DNA sequence is further from the island of good sequences. For this reason I conjecture that, at least early in evolution, and perhaps still now, the genetic code was not implemented perfectly but was implemented noisily, with some codons coding for a distribution of possible amino acids. This noisy code could even be switched on and off from cell to cell in an organism by having multiple aminoacyl-tRNA synthetases, some more reliable than others. Whilst our model assumed that the bits of the genome do not interact, ignored the fact that the information is represented redundantly, assumed that there is a direct relationship between phenotypic fitness and the genotype, and assumed that the crossover probability in recombination is high, I believe these qualitative results would still hold if more complex models of √ fitness and crossover were used: the relative benefit of sex will still scale as G. Only in small, in-bred populations are the benefits of sex expected to be diminished. In summary: Why have sex? Because sex is good for your bits!

Further reading How did a high-information-content self-replicating system ever emerge in the first place? In the general area of the origins of life and other tricky questions about evolution, I highly recommend Maynard Smith and Sz´athmary (1995), Maynard Smith and Sz´athmary (1999), Kondrashov (1988), Maynard Smith (1988), Ridley (2000), Dyson (1985), Cairns-Smith (1985), and Hopfield (1978).

19.6 Further exercises Exercise 19.4. [3 ] How good must the error-correcting machinery in DNA replication be, given that mammals have not all died out long ago? Estimate the probability of nucleotide substitution, per cell division. [See Appendix C.4.]

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280

19 — Why have Sex? Information Acquisition and Evolution

Exercise 19.5. [4 ] Given that DNA replication is achieved by bumbling Brownian motion and ordinary thermodynamics in a biochemical porridge at a temperature of 35 C, it’s astonishing that the error-rate of DNA replication is about 10−9 per replicated nucleotide. How can this reliability be achieved, given that the energetic difference between a correct base-pairing and an incorrect one is only one or two hydrogen bonds and the thermal energy kT is only about a factor of four smaller than the free energy associated with a hydrogen bond? If ordinary thermodynamics is what favours correct base-pairing, surely the frequency of incorrect base-pairing should be about f = exp(−∆E/kT ),

(19.21)

where ∆E is the free energy difference, i.e., an error frequency of f ' 10 −4 ? How has DNA replication cheated thermodynamics? The situation is equally perplexing in the case of protein synthesis, which translates an mRNA sequence into a polypeptide in accordance with the genetic code. Two specific chemical reactions are protected against errors: the binding of tRNA molecules to amino acids, and the production of the polypeptide in the ribosome, which, like DNA replication, involves base-pairing. Again, the fidelity is high (an error rate of about 10 −4 ), and this fidelity can’t be caused by the energy of the ‘correct’ final state being especially low – the correct polypeptide sequence is not expected to be significantly lower in energy than any other sequence. How do cells perform error correction? (See Hopfield (1974), Hopfield (1980)). Exercise 19.6. [2 ] While the genome acquires information through natural selection at a rate of a few bits per generation, your brain acquires information at a greater rate. Estimate at what rate new information can be stored in long term memory by your brain. Think of learning the words of a new language, for example.

19.7 Solutions Solution to exercise 19.1 (p.275). For small enough N , whilst the average fitness of the population increases, some unlucky bits become frozen into the bad state. (These bad genes are sometimes known as hitchhikers.) The homogeneity assumption breaks down. Eventually, all individuals have identical genotypes that are mainly 1-bits, but contain some 0-bits too. The smaller the population, the greater the number of frozen 0-bits is expected to be. How small can the population size N be if the theory of sex is accurate? We find experimentally that the theory based on assuming homogeneity fits √ poorly only if√ the population size N is smaller than ∼ G. If N is significantly smaller than G, information cannot possibly be acquired at a rate as big as √ G, since the information content of the Blind Watchmaker’s decisions cannot be any greater than 2N bits per generation, this being the number of bits required to specify which of the 2N children get to reproduce. Baum et al. (1995),√analyzing a similar model, show that the population size N should be about G(log G)2 to make hitchhikers unlikely to arise.

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Part IV

Probabilities and Inference

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About Part IV The number of inference problems that can (and perhaps should) be tackled by Bayesian inference methods is enormous. In this book, for example, we discuss the decoding problem for error-correcting codes, the task of inferring clusters from data, the task of interpolation through noisy data, and the task of classifying patterns given labelled examples. Most techniques for solving these problems can be categorized as follows. Exact methods compute the required quantities directly. Only a few interesting problems have a direct solution, but exact methods are important as tools for solving subtasks within larger problems. Methods for the exact solution of inference problems are the subject of Chapters 21, 24, 25, and 26. Approximate methods can be subdivided into 1. deterministic approximations, which include maximum likelihood (Chapter 22), Laplace’s method (Chapters 27 and 28) and variational methods (Chapter 33); and 2. Monte Carlo methods – techniques in which random numbers play an integral part – which will be discussed in Chapters 29, 30, and 32. This part of the book does not form a one-dimensional story. Rather, the ideas make up a web of interrelated threads which will recombine in subsequent chapters. Chapter 3, which is an honorary member of this part, discussed a range of simple examples of inference problems and their Bayesian solutions. To give further motivation for the toolbox of inference methods discussed in this part, Chapter 20 discusses the problem of clustering; subsequent chapters discuss the probabilistic interpretation of clustering as mixture modelling. Chapter 21 discusses the option of dealing with probability distributions by completely enumerating all hypotheses. Chapter 22 introduces the idea of maximization methods as a way of avoiding the large cost associated with complete enumeration, and points out reasons why maximum likelihood is not good enough. Chapter 23 reviews the probability distributions that arise most often in Bayesian inference. Chapters 24, 25, and 26 discuss another way of avoiding the cost of complete enumeration: marginalization. Chapter 25 discusses message-passing methods appropriate for graphical models, using the decoding of error-correcting codes as an example. Chapter 26 combines these ideas with message-passing concepts from Chapters 16 and 17. These chapters are a prerequisite for the understanding of advanced error-correcting codes. Chapter 27 discusses deterministic approximations including Laplace’s method. This chapter is a prerequisite for understanding the topic of complexity control in learning algorithms, an idea that is discussed in general terms in Chapter 28. 282

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About Part IV Chapter 29 discusses Monte Carlo methods. Chapter 30 gives details of state-of-the-art Monte Carlo techniques. Chapter 31 introduces the Ising model as a test-bed for probabilistic methods. An exact message-passing method and a Monte Carlo method are demonstrated. A motivation for studying the Ising model is that it is intimately related to several neural network models. Chapter 32 describes ‘exact’ Monte Carlo methods and demonstrates their application to the Ising model. Chapter 33 discusses variational methods and their application to Ising models and to simple statistical inference problems including clustering. This chapter will help the reader understand the Hopfield network (Chapter 42) and the EM algorithm, which is an important method in latent-variable modelling. Chapter 34 discusses a particularly simple latent variable model called independent component analysis. Chapter 35 discusses a ragbag of assorted inference topics. Chapter 36 discusses a simple example of decision theory. Chapter 37 discusses differences between sampling theory and Bayesian methods.

A theme: what inference is about A widespread misconception is that the aim of inference is to find the most probable explanation for some data. While this most probable hypothesis may be of interest, and some inference methods do locate it, this hypothesis is just the peak of a probability distribution, and it is the whole distribution that is of interest. As we saw in Chapter 4, the most probable outcome from a source is often not a typical outcome from that source. Similarly, the most probable hypothesis given some data may be atypical of the whole set of reasonablyplausible hypotheses.

About Chapter 20 Before reading the next chapter, exercise 2.17 (p.36) and section 11.2 (inferring the input to a Gaussian channel) are recommended reading.

283

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20 An Example Inference Task: Clustering

Human brains are good at finding regularities in data. One way of expressing regularity is to put a set of objects into groups that are similar to each other. For example, biologists have found that most objects in the natural world fall into one of two categories: things that are brown and run away, and things that are green and don’t run away. The first group they call animals, and the second, plants. We’ll call this operation of grouping things together clustering. If the biologist further sub-divides the cluster of plants into subclusters, we would call this ‘hierarchical clustering’; but we won’t be talking about hierarchical clustering yet. In this chapter we’ll just discuss ways to take a set of N objects and group them into K clusters. There are several motivations for clustering. First, a good clustering has predictive power. When an early biologist encounters a new green thing he has not seen before, his internal model of plants and animals fills in predictions for attributes of the green thing: it’s unlikely to jump on him and eat him; if he touches it, he might get grazed or stung; if he eats it, he might feel sick. All of these predictions, while uncertain, are useful, because they help the biologist invest his resources (for example, the time spent watching for predators) well. Thus, we perform clustering because we believe the underlying cluster labels are meaningful, will lead to a more efficient description of our data, and will help us choose better actions. This type of clustering is sometimes called ‘mixture density modelling’, and the objective function that measures how well the predictive model is working is the information content of the data, log 1/P ({x}). Second, clusters can be a useful aid to communication because they allow lossy compression. The biologist can give directions to a friend such as ‘go to the third tree on the right then take a right turn’ (rather than ‘go past the large green thing with red berries, then past the large green thing with thorns, then . . .’). The brief category name ‘tree’ is helpful because it is sufficient to identify an object. Similarly, in lossy image compression, the aim is to convey in as few bits as possible a reasonable reproduction of a picture; one way to do this is to divide the image into N small patches, and find a close match to each patch in an alphabet of K image-templates; then we send a close fit to the image by sending the list of labels k 1 , k2 , . . . , kN of the matching templates. The task of creating a good library of image-templates is equivalent to finding a set of cluster centres. This type of clustering is sometimes called ‘vector quantization’. We can formalize a vector quantizer in terms of an assignment rule x → k(x) for assigning datapoints x to one of K codenames, and a reconstruction rule k → m(k) , the aim being to choose the functions k(x) and m (k) so as to 284

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20.1: K-means clustering

285

minimize the expected distortion, which might be defined to be D=

X

P (x)

x

i2 1 h (k(x)) m −x . 2

(20.1)

[The ideal objective function would be to minimize the psychologically perceived distortion of the image. Since it is hard to quantify the distortion perceived by a human, vector quantization and lossy compression are not so crisply defined problems as data modelling and lossless compression.] In vector quantization, we don’t necessarily believe that the templates {m (k) } have any natural meaning; they are simply tools to do a job. We note in passing the similarity of the assignment rule (i.e., the encoder) of vector quantization to the decoding problem when decoding an error-correcting code. A third reason for making a cluster model is that failures of the cluster model may highlight interesting objects that deserve special attention. If we have trained a vector quantizer to do a good job of compressing satellite pictures of ocean surfaces, then maybe patches of image that are not well compressed by the vector quantizer are the patches that contain ships! If the biologist encounters a green thing and sees it run (or slither) away, this misfit with his cluster model (which says green things don’t run away) cues him to pay special attention. One can’t spend all one’s time being fascinated by things; the cluster model can help sift out from the multitude of objects in one’s world the ones that really deserve attention. A fourth reason for liking clustering algorithms is that they may serve as models of learning processes in neural systems. The clustering algorithm that we now discuss, the K-means algorithm, is an example of a competitive learning algorithm. The algorithm works by having the K clusters compete with each other for the right to own the data points.

Figure 20.1. N = 40 data points.

20.1 K-means clustering The K-means algorithm is an algorithm for putting N data points in an Idimensional space into K clusters. Each cluster is parameterized by a vector m(k) called its mean. The data points will be denoted by {x (n) } where the superscript n runs from 1 to the number of data points N . Each vector x has I components x i . We will assume that the space that x lives in is a real space and that we have a metric that defines distances between points, for example, d(x, y) =

1X (xi − yi )2 . 2

(20.2)

i

To start the K-means algorithm (algorithm 20.2), the K means {m (k) } are initialized in some way, for example to random values. K-means is then an iterative two-step algorithm. In the assignment step, each data point n is assigned to the nearest mean. In the update step, the means are adjusted to match the sample means of the data points that they are responsible for. The K-means algorithm is demonstrated for a toy two-dimensional data set in figure 20.3, where 2 means are used. The assignments of the points to the two clusters are indicated by two point styles, and the two means are shown by the circles. The algorithm converges after three iterations, at which point the assignments are unchanged so the means remain unmoved when updated. The K-means algorithm always converges to a fixed point.

About the name... As far as I know, the ‘K’ in K-means clustering simply refers to the chosen number of clusters. If Newton had followed the same naming policy, maybe we would learn at school about ‘calculus for the variable x’. It’s a silly name, but we are stuck with it.

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20 — An Example Inference Task: Clustering

Algorithm 20.2. The K-means clustering algorithm.

Initialization. Set K means {m(k) } to random values. Assignment step. Each data point n is assigned to the nearest mean. We denote our guess for the cluster k (n) that the point x(n) belongs to by kˆ(n) . kˆ(n) = argmin{d(m(k) , x(n) )}. (20.3) k

An alternative, equivalent representation of this assignment of points to clusters is given by ‘responsibilities’, which are indicator (n) (n) variables rk . In the assignment step, we set rk to one if mean k (n) is the closest mean to datapoint x(n) ; otherwise rk is zero. (n)

rk

=



1 0

kˆ(n) = k kˆ (n) 6= k.

if if

(20.4)

What about ties? – We don’t expect two means to be exactly the same distance from a data point, but if a tie does happen, kˆ (n) is set to the smallest of the winning {k}. Update step. The model parameters, the means, are adjusted to match the sample means of the data points that they are responsible for. X (n) rk x(n) m(k) =

n

(20.5)

R(k)

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(20.6)

n

What about means with no responsibilities? – If R (k) = 0, then we leave the mean m(k) where it is. Repeat the assignment step and update step until ments do not change.

the

assign-

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20.1: K-means clustering

287 Figure 20.3. K-means algorithm applied to a data set of 40 points. K = 2 means evolve to stable locations after three iterations.

Data:

Assignment

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Exercise 20.1.[4, p.291] See if you can prove that K-means always converges. [Hint: find a physical analogy and an associated Lyapunov function.] [A Lyapunov function is a function of the state of the algorithm that decreases whenever the state changes and that is bounded below. If a system has a Lyapunov function then its dynamics converge.] The K-means algorithm with a larger number of means, 4, is demonstrated in figure 20.4. The outcome of the algorithm depends on the initial condition. In the first case, after five iterations, a steady state is found in which the data points are fairly evenly split between the four clusters. In the second case, after six iterations, half the data points are in one cluster, and the others are shared among the other three clusters.

Questions about this algorithm The K-means algorithm has several ad hoc features. Why does the update step set the ‘mean’ to the mean of the assigned points? Where did the distance d come from? What if we used a different measure of distance between x and m? How can we choose the ‘best’ distance? [In vector quantization, the distance

Assignment

Update

Figure 20.4. K-means algorithm applied to a data set of 40 points. Two separate runs, both with K = 4 means, reach different solutions. Each frame shows a successive assignment step.

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288

20 — An Example Inference Task: Clustering 10

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Figure 20.5. K-means algorithm for a case with two dissimilar clusters. (a) The “little ’n’ large” data. (b) A stable set of assignments and means. Note that four points belonging to the broad cluster have been incorrectly assigned to the narrower cluster. (Points assigned to the right-hand cluster are shown by plus signs.)

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Figure 20.6. Two elongated clusters, and the stable solution found by the K-means algorithm. (a)

(b)

function is provided as part of the problem definition; but I’m assuming we are interested in data-modelling rather than vector quantization.] How do we choose K? Having found multiple alternative clusterings for a given K, how can we choose among them?

Cases where K-means might be viewed as failing. Further questions arise when we look for cases where the algorithm behaves badly (compared with what the man in the street would call ‘clustering’). Figure 20.5a shows a set of 75 data points generated from a mixture of two Gaussians. The right-hand Gaussian has less weight (only one fifth of the data points), and it is a less broad cluster. Figure 20.5b shows the outcome of using K-means clustering with K = 2 means. Four of the big cluster’s data points have been assigned to the small cluster, and both means end up displaced to the left of the true centres of the clusters. The K-means algorithm takes account only of the distance between the means and the data points; it has no representation of the weight or breadth of each cluster. Consequently, data points that actually belong to the broad cluster are incorrectly assigned to the narrow cluster. Figure 20.6 shows another case of K-means behaving badly. The data evidently fall into two elongated clusters. But the only stable state of the K-means algorithm is that shown in figure 20.6b: the two clusters have been sliced in half! These two examples show that there is something wrong with the distance d in the K-means algorithm. The K-means algorithm has no way of representing the size or shape of a cluster. A final criticism of K-means is that it is a ‘hard’ rather than a ‘soft’ algorithm: points are assigned to exactly one cluster and all points assigned to a cluster are equals in that cluster. Points located near the border between two or more clusters should, arguably, play a partial role in determining the locations of all the clusters that they could plausibly be assigned to. But in the K-means algorithm, each borderline point is dumped in one cluster, and

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20.2: Soft K-means clustering

289

has an equal vote with all the other points in that cluster, and no vote in any other clusters.

20.2 Soft K-means clustering These criticisms of K-means motivate the ‘soft K-means algorithm’, algorithm 20.7. The algorithm has one parameter, β, which we could term the stiffness. Assignment step. Each data point x(n) is given a soft ‘degree of assignment’ to each of the means. We call the degree to which x (n) (n) is assigned to cluster k the responsibility r k (the responsibility of cluster k for point n).  exp −β d(m(k) , x(n) ) (n) . P (20.7) rk = (k 0 ) , x(n) ) k 0 exp −β d(m The sum of the K responsibilities for the nth point is 1.

Update step. The model parameters, the means, are adjusted to match the sample means of the data points that they are responsible for. X (n) rk x(n) m(k) =

n

R(k)

where R(k) is the total responsibility of mean k, X (n) rk . R(k) =

(20.8)

(20.9)

n

Notice the similarity of this soft K-means algorithm to the hard K-means algorithm 20.2. The update step is identical; the only difference is that the (n) responsibilities rk can take on values between 0 and 1. Whereas the assignment kˆ (n) in the K-means algorithm involved a ‘min’ over the distances, the rule for assigning the responsibilities is a ‘soft-min’ (20.7). . Exercise 20.2.[2 ] Show that as the stiffness β goes to ∞, the soft K-means algorithm becomes identical to the original hard K-means algorithm, except for the way in which means with no assigned points behave. Describe what those means do instead of sitting still. Dimensionally, the stiffness√β is an inverse-length-squared, so we can associate a lengthscale, σ ≡ 1/ β, with it. The soft K-means algorithm is demonstrated in figure 20.8. The lengthscale is shown by the radius of the circles surrounding the four means. Each panel shows the final fixed point reached for a different value of the lengthscale σ.

20.3 Conclusion At this point, we may have fixed some of the problems with the original Kmeans algorithm by introducing an extra complexity-control parameter β. But how should we set β? And what about the problem of the elongated clusters,

Algorithm 20.7. Soft K-means algorithm, version 1.

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290

20 — An Example Inference Task: Clustering

Large σ . . .

...

. . . small σ

Figure 20.8. Soft K-means algorithm, version 1, applied to a data set of 40 points. K = 4. Implicit lengthscale parameter σ = 1/β 1/2 varied from a large to a small value. Each picture shows the state of all four means, with the implicit lengthscale shown by the radius of the four circles, after running the algorithm for several tens of iterations. At the largest lengthscale, all four means converge exactly to the data mean. Then the four means separate into two groups of two. At shorter lengthscales, each of these pairs itself bifurcates into subgroups.

and the clusters of unequal weight and width? Adding one stiffness parameter β is not going to make all these problems go away. We’ll come back to these questions in a later chapter, as we develop the mixture-density-modelling view of clustering.

Further reading For a vector-quantization approach to clustering see (Luttrell, 1989; Luttrell, 1990).

20.4 Exercises . Exercise 20.3.[3, p.291] Explore the properties of the soft K-means algorithm, version 1, assuming that the datapoints {x} come from a single separable two-dimensional Gaussian distribution with mean zero and variances (var(x1 ), var(x2 )) = (σ12 , σ22 ), with σ12 > σ22 . Set K = 2, assume N is large, and investigate the fixed points of the algorithm as β is varied. [Hint: assume that m(1) = (m, 0) and m(2) = (−m, 0).]

. Exercise 20.4.[3 ] Consider the soft K-means algorithm applied to a large amount of one-dimensional data that comes from a mixture of two equalweight Gaussians with true means µ = ±1 and standard deviation σ P , for example σP = 1. Show that the hard K-means algorithm with K = 2 leads to a solution in which the two means are further apart than the two true means. Discuss what happens for other values of β, and find the value of β such that the soft algorithm puts the two means in the correct places.

-1

1

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20.5: Solutions

291

20.5 Solutions Solution to exercise 20.1 (p.287). We can associate an ‘energy’ with the state of the K-means algorithm by connecting a spring between each point x (n) and the mean that is responsible for it. The energy of one spring is proportional to its squared length, namely βd(x(n) , m(k) ) where β is the stiffness of the spring. The total energy of all the springs is a Lyapunov function for the algorithm, because (a) the assignment step can only decrease the energy – a point only changes its allegiance if the length of its spring would be reduced; (b) the update step can only decrease the energy – moving m (k) to the mean is the way to minimize the energy of its springs; and (c) the energy is bounded below – which is the second condition for a Lyapunov function. Since the algorithm has a Lyapunov function, it converges.

m1 m2

Solution to exercise 20.3 (p.290). If the means are initialized to m (1) = (m, 0) and m(1) = (−m, 0), the assignment step for a point at location x 1 , x2 gives exp(−β(x1 − m)2 /2) exp(−β(x1 − m)2 /2) + exp(−β(x1 + m)2 /2) 1 , 1 + exp(−2βmx1 )

r1 (x) = =

m

R

dx P (x1 ) x1 r1 (x) R 1 dx1 P (x1 ) r1 (x) Z 1 = 2 dx1 P (x1 ) x1 . 1 + exp(−2βmx1 )

=

(20.11)

Figure 20.9. Schematic diagram of the bifurcation as the largest data variance σ1 increases from below 1/β 1/2 to above 1/β 1/2 . The data variance is indicated by the ellipse.

(20.12) (20.13)

Now, m = 0 is a fixed point, but the question is, is it stable or unstable? For tiny m (that is, βσ1 m  1), we can Taylor-expand 1 1 ' (1 + βmx1 ) + · · · 1 + exp(−2βmx1 ) 2

m1

(20.10)

and the updated m is 0

m2

(20.14)

4 Data density Mean locations

3 2 1

-2-1 0 1 2

0 -1 -2-1 0 1 2

-2

so

-3

m0 '

Z

-4

dx1 P (x1 ) x1 (1 + βmx1 )

= σ12 βm.

0

(20.15) (20.16)

For small m, m either grows or decays exponentially under this mapping, depending on whether σ12 β is greater than or less than 1. The fixed point m = 0 is stable if σ12 ≤ 1/β (20.17) and unstable otherwise. [Incidentally, this derivation shows that this result is general, holding for any true probability distribution P (x 1 ) having variance σ12 , not just the Gaussian.] If σ12 > 1/β then there is a bifurcation and there are two stable fixed points surrounding the unstable fixed point at m = 0. To illustrate this bifurcation, figure 20.10 shows the outcome of running the soft K-means algorithm with β = 1 on one-dimensional data with standard deviation σ 1 for various values of σ1 . Figure 20.11 shows this pitchfork bifurcation from the other point of view, where the data’s standard deviation σ 1 is fixed and the algorithm’s lengthscale σ = 1/β 1/2 is varied on the horizontal axis.

0.5

1

1.5

2

2.5

3

3.5

4

Figure 20.10. The stable mean locations as a function of σ1 , for constant β, found numerically (thick lines), and the approximation (20.22) (thin lines).

0.8 Data density Mean locns.

0.6 0.4 0.2

-2 -1 0 1 2

0 -0.2 -2 -1 0 1 2

-0.4 -0.6 -0.8 0

0.5

1

1.5

2

Figure 20.11. The stable mean locations as a function of 1/β 1/2 , for constant σ1 .

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292

20 — An Example Inference Task: Clustering Here is a cheap theory to model how the fitted parameters ±m behave beyond the bifurcation, based on continuing the series expansion. This continuation of the series is rather suspect, since the series isn’t necessarily expected to converge beyond the bifurcation point, but the theory fits well anyway. We take our analytic approach one term further in the expansion 1 1 1 ' (1 + βmx1 − (βmx1 )3 ) + · · · 1 + exp(−2βmx1 ) 2 3

(20.18)

then we can solve for the shape of the bifurcation to leading order, which depends on the fourth moment of the distribution: Z 1 m0 ' dx1 P (x1 )x1 (1 + βmx1 − (βmx1 )3 ) (20.19) 3 1 (20.20) = σ12 βm − (βm)3 3σ14 . 3 [At (20.20) we use the fact that P (x1 ) is Gaussian to find the fourth moment.] This map has a fixed point at m such that σ12 β(1 − (βm)2 σ12 ) = 1,

(20.21)

i.e., m = ±β −1/2

(σ12 β − 1)1/2 . σ12 β

(20.22)

The thin line in figure 20.10 shows this theoretical approximation. Figure 20.10 shows the bifurcation as a function of σ1 for fixed β; figure 20.11 shows the bifurcation as a function of 1/β 1/2 for fixed σ1 .

. Exercise 20.5.[2, p.292] Why does the pitchfork in figure 20.11 tend to the values ∼ ±0.8 as 1/β 1/2 → 0? Give an analytic expression for this asymptote. Solution to exercise 20.5 (p.292). The asymptote is the mean of the rectified Gaussian, R∞ p 0 Normal(x, 1)x dx = 2/π ' 0.798. (20.23) 1/2

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21 Exact Inference by Complete Enumeration We open our toolbox of methods for handling probabilities by discussing a brute-force inference method: complete enumeration of all hypotheses, and evaluation of their probabilities. This approach is an exact method, and the difficulty of carrying it out will motivate the smarter exact and approximate methods introduced in the following chapters.

21.1 The burglar alarm Bayesian probability theory is sometimes called ‘common sense, amplified’. When thinking about the following questions, please ask your common sense what it thinks the answers are; we will then see how Bayesian methods confirm your everyday intuition. Example 21.1. Fred lives in Los Angeles and commutes 60 miles to work. Whilst at work, he receives a phone-call from his neighbour saying that Fred’s burglar alarm is ringing. What is the probability that there was a burglar in his house today? While driving home to investigate, Fred hears on the radio that there was a small earthquake that day near his home. ‘Oh’, he says, feeling relieved, ‘it was probably the earthquake that set off the alarm’. What is the probability that there was a burglar in his house? (After Pearl, 1988). Let’s introduce variables b (a burglar was present in Fred’s house today), a (the alarm is ringing), p (Fred receives a phonecall from the neighbour reporting the alarm), e (a small earthquake took place today near Fred’s house), and r (the radio report of earthquake is heard by Fred). The probability of all these variables might factorize as follows: P (b, e, a, p, r) = P (b)P (e)P (a | b, e)P (p | a)P (r | e),

(21.1)

and plausible values for the probabilities are: 1. Burglar probability: P (b = 1) = β, P (b = 0) = 1 − β,

(21.2)

e.g., β = 0.001 gives a mean burglary rate of once every three years. 2. Earthquake probability: P (e = 1) = , P (e = 0) = 1 − , 293

(21.3)

Earthquake Burglar j j

j

Radio

@ R @

jAlarm @ R @ j

Phonecall

Figure 21.1. Belief network for the burglar alarm problem.

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294

21 — Exact Inference by Complete Enumeration with, e.g.,  = 0.001; our assertion that the earthquakes are independent of burglars, i.e., the prior probability of b and e is P (b, e) = P (b)P (e), seems reasonable unless we take into account opportunistic burglars who strike immediately after earthquakes.

3. Alarm ringing probability: we assume the alarm will ring if any of the following three events happens: (a) a burglar enters the house, and triggers the alarm (let’s assume the alarm has a reliability of α b = 0.99, i.e., 99% of burglars trigger the alarm); (b) an earthquake takes place, and triggers the alarm (perhaps αe = 1% of alarms are triggered by earthquakes?); or (c) some other event causes a false alarm; let’s assume the false alarm rate f is 0.001, so Fred has false alarms from non-earthquake causes once every three years. [This type of dependence of a on b and e is known as a ‘noisy-or’.] The probabilities of a given b and e are then: P (a = 0 | b = 0, P (a = 0 | b = 1, P (a = 0 | b = 0, P (a = 0 | b = 1,

e = 0) e = 0) e = 1) e = 1)

= = = =

(1 − f ), (1 − f )(1 − αb ), (1 − f )(1 − αe ), (1 − f )(1 − αb )(1 − αe ),

e = 0) e = 0) e = 1) e = 1)

= = = =

0.999, 0.009 99, 0.989 01, 0.009 890 1,

P (a = 1 | b = 0, P (a = 1 | b = 1, P (a = 1 | b = 0, P (a = 1 | b = 1,

e = 0) e = 0) e = 1) e = 1)

= = = =

f 1 − (1 − f )(1 − αb ) 1 − (1 − f )(1 − αe ) 1 − (1 − f )(1 − αb )(1 − αe )

or, in numbers, P (a = 0 | b = 0, P (a = 0 | b = 1, P (a = 0 | b = 0, P (a = 0 | b = 1,

P (a = 1 | b = 0, P (a = 1 | b = 1, P (a = 1 | b = 0, P (a = 1 | b = 1,

e = 0) e = 0) e = 1) e = 1)

= = = =

0.001 0.990 01 0.010 99 0.990 109 9.

We assume the neighbour would never phone if the alarm is not ringing [P (p = 1 | a = 0) = 0]; and that the radio is a trustworthy reporter too [P (r = 1 | e = 0) = 0]; we won’t need to specify the probabilities P (p = 1 | a = 1) or P (r = 1 | e = 1) in order to answer the questions above, since the outcomes p = 1 and r = 1 give us certainty respectively that a = 1 and e = 1. We can answer the two questions about the burglar by computing the posterior probabilities of all hypotheses given the available information. Let’s start by reminding ourselves that the probability that there is a burglar, before either p or r is observed, is P (b = 1) = β = 0.001, and the probability that an earthquake took place is P (e = 1) =  = 0.001, and these two propositions are independent. First, when p = 1, we know that the alarm is ringing: a = 1. The posterior probability of b and e becomes: P (b, e | a = 1) =

P (a = 1 | b, e)P (b)P (e) . P (a = 1)

(21.4)

The numerator’s four possible values are P (a = 1 | b = 0, P (a = 1 | b = 1, P (a = 1 | b = 0, P (a = 1 | b = 1,

e = 0) × P (b = 0) × P (e = 0) e = 0) × P (b = 1) × P (e = 0) e = 1) × P (b = 0) × P (e = 1) e = 1) × P (b = 1) × P (e = 1)

= = = =

0.001 × 0.999 × 0.999 0.990 01 × 0.001 × 0.999 0.010 99 × 0.999 × 0.001 0.990 109 9 × 0.001 × 0.001

The normalizing constant is the sum of these four numbers, P (a = 1) = 0.002, and the posterior probabilities are P (b = 0, P (b = 1, P (b = 0, P (b = 1,

e = 0 | a = 1) e = 0 | a = 1) e = 1 | a = 1) e = 1 | a = 1)

= = = =

0.4993 0.4947 0.0055 0.0005.

(21.5)

= = = =

0.000 998 0.000 989 0.000 010 979 9.9 × 10 −7 .

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21.2: Exact inference for continuous hypothesis spaces

295

To answer the question, ‘what’s the probability a burglar was there?’ we marginalize over the earthquake variable e: P (b = 0 | a = 1) = P (b = 0, e = 0 | a = 1) + P (b = 0, e = 1 | a = 1) = 0.505 P (b = 1 | a = 1) = P (b = 1, e = 0 | a = 1) + P (b = 1, e = 1 | a = 1) = 0.495. (21.6) So there is nearly a 50% chance that there was a burglar present. It is important to note that the variables b and e, which were independent a priori, are now dependent. The posterior distribution (21.5) is not a separable function of b and e. This fact is illustrated most simply by studying the effect of learning that e = 1. When we learn e = 1, the posterior probability of b is given by P (b | e = 1, a = 1) = P (b, e = 1 | a = 1)/P (e = 1 | a = 1), i.e., by dividing the bottom two rows of (21.5), by their sum P (e = 1 | a = 1) = 0.0060. The posterior probability of b is: P (b = 0 | e = 1, a = 1) = 0.92 P (b = 1 | e = 1, a = 1) = 0.08.

(21.7)

There is thus now an 8% chance that a burglar was in Fred’s house. It is in accordance with everyday intuition that the probability that b = 1 (a possible cause of the alarm) reduces when Fred learns that an earthquake, an alternative explanation of the alarm, has happened.

Explaining away This phenomenon, that one of the possible causes (b = 1) of some data (the data in this case being a = 1) becomes less probable when another of the causes (e = 1) becomes more probable, even though those two causes were independent variables a priori, is known as explaining away. Explaining away is an important feature of correct inferences, and one that any artificial intelligence should replicate. If we believe that the neighbour and the radio service are unreliable or capricious, so that we are not certain that the alarm really is ringing or that an earthquake really has happened, the calculations become more complex, but the explaining-away effect persists; the arrival of the earthquake report r simultaneously makes it more probable that the alarm truly is ringing, and less probable that the burglar was present. In summary, we solved the inference questions about the burglar by enumerating all four hypotheses about the variables (b, e), finding their posterior probabilities, and marginalizing to obtain the required inferences about b. . Exercise 21.2.[2 ] After Fred receives the phone-call about the burglar alarm, but before he hears the radio report, what, from his point of view, is the probability that there was a small earthquake today?

21.2 Exact inference for continuous hypothesis spaces Many of the hypothesis spaces we will consider are naturally thought of as continuous. For example, the unknown decay length λ of section 3.1 (p.48) lives in a continuous one-dimensional space; and the unknown mean and standard deviation of a Gaussian µ, σ live in a continuous two-dimensional space. In any practical computer implementation, such continuous spaces will necessarily be discretized, however, and so can, in principle, be enumerated – at a grid of parameter values, for example. In figure 3.2 we plotted the likelihood

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296

21 — Exact Inference by Complete Enumeration Figure 21.2. Enumeration of an entire (discretized) hypothesis space for one Gaussian with parameters µ (horizontal axis) and σ (vertical).

function for the decay length as a function of λ by evaluating the likelihood at a finely-spaced series of points.

A two-parameter model Let’s look at the Gaussian distribution as an example of a model with a twodimensional hypothesis space. The one-dimensional Gaussian distribution is parameterized by a mean µ and a standard deviation σ:   1 (x − µ)2 ≡ Normal(x; µ, σ 2 ). (21.8) P (x | µ, σ) = √ exp − 2σ 2 2πσ Figure 21.2 shows an enumeration of one hundred hypotheses about the mean and standard deviation of a one-dimensional Gaussian distribution. These hypotheses are evenly spaced in a ten by ten square grid covering ten values of µ and ten values of σ. Each hypothesis is represented by a picture showing the probability density that it puts on x. We now examine the inference of µ and σ given data points xn , n = 1, . . . , N , assumed to be drawn independently from this density. Imagine that we acquire data, for example the five points shown in figure 21.3. We can now evaluate the posterior probability of each of the one hundred subhypotheses by evaluating the likelihood of each, that is, the value of P ({xn }5n=1 | µ, σ). The likelihood values are shown diagrammatically in figure 21.4 using the line thickness to encode the value of the likelihood. Subhypotheses with likelihood smaller than e −8 times the maximum likelihood have been deleted. Using a finer grid, we can represent the same information by plotting the likelihood as a surface plot or contour plot as a function of µ and σ (figure 21.5).

A five-parameter mixture model Eyeballing the data (figure 21.3), you might agree that it seems more plausible that they come not from a single Gaussian but from a mixture of two Gaussians, defined by two means, two standard deviations, and two mixing

-0.5

0

0.5

1

1.5

2

2.5

Figure 21.3. Five datapoints {xn }5n=1 . The horizontal coordinate is the value of the datum, xn ; the vertical coordinate has no meaning.

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21.2: Exact inference for continuous hypothesis spaces

297

Figure 21.4. Likelihood function, given the data of figure 21.3, represented by line thickness. Subhypotheses having likelihood smaller than e−8 times the maximum likelihood are not shown.

0.06 1 0.05

0.9

0.04

0.8

0.03

0.7 0.6

0.02

0.5 0.01

0.4

10

0.3 0.2

0.8 0.6 sigma

0.1 0

0.4 0.2 0

0.5

1 mean

1.5

0.5

1

2 mean

1.5

2

sigma

Figure 21.5. The likelihood function for the parameters of a Gaussian distribution. Surface plot and contour plot of the log likelihood as a function of µ and σ. The data set of N = 5 points x ¯ = 1.0 and Phad mean S = (x − x ¯)2 = 1.0.

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298

21 — Exact Inference by Complete Enumeration Figure 21.6. Enumeration of the entire (discretized) hypothesis space for a mixture of two Gaussians. Weight of the mixture components is π1 , π2 = 0.6, 0.4 in the top half and 0.8, 0.2 in the bottom half. Means µ1 and µ2 vary horizontally, and standard deviations σ1 and σ2 vary vertically.

coefficients π1 and π2 , satisfying π1 + π2 = 1, πi ≥ 0.     2 2 π2 π1 1) 2) +√ P (x | µ1 , σ1 , π1 , µ2 , σ2 , π2 ) = √ exp − (x−µ exp − (x−µ 2 2 2σ 2σ 1 2 2πσ1 2πσ2 Let’s enumerate the subhypotheses for this alternative model. The parameter space is five-dimensional, so it becomes challenging to represent it on a single page. Figure 21.6 enumerates 800 subhypotheses with different values of the five parameters µ1 , µ2 , σ1 , σ2 , π1 . The means are varied between five values each in the horizontal directions. The standard deviations take on four values each vertically. And π1 takes on two values vertically. We can represent the inference about these five parameters in the light of the five datapoints as shown in figure 21.7. If we wish to compare the one-Gaussian model with the mixture-of-two model, we can find the models’ posterior probabilities by evaluating the marginal likelihood or evidence for each model H, P ({x} | H). The evidence is given by integrating over the parameters, θ; the integration can be implemented numerically by summing over the alternative enumerated values of θ, X P ({x} | H) = P (θ)P ({x} | θ, H), (21.9) where P (θ) is the prior distribution over the grid of parameter values, which I take to be uniform. For the mixture of two Gaussians this integral is a five-dimensional integral; if it is to be performed at all accurately, the grid of points will need to be much finer than the grids shown in the figures. If the uncertainty about each of K parameters has been reduced by, say, a factor of ten by observing the data, then brute-force integration requires a grid of at least 10 K points. This

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21.2: Exact inference for continuous hypothesis spaces

299 Figure 21.7. Inferring a mixture of two Gaussians. Likelihood function, given the data of figure 21.3, represented by line thickness. The hypothesis space is identical to that shown in figure 21.6. Subhypotheses having likelihood smaller than e−8 times the maximum likelihood are not shown, hence the blank regions, which correspond to hypotheses that the data have ruled out.

-0.5

exponential growth of computation with model size is the reason why complete enumeration is rarely a feasible computational strategy. Exercise 21.3.[1 ] Imagine fitting a mixture of ten Gaussians to data in a twenty-dimensional space. Estimate the computational cost of implementing inferences for this model by enumeration of a grid of parameter values.

0

0.5

1

1.5

2

2.5

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22 Maximum Likelihood and Clustering Rather than enumerate all hypotheses – which may be exponential in number – we can save a lot of time by homing in on one good hypothesis that fits the data well. This is the philosophy behind the maximum likelihood method, which identifies the setting of the parameter vector θ that maximizes the likelihood, P (Data | θ, H). For some models the maximum likelihood parameters can be identified instantly from the data; for more complex models, finding the maximum likelihood parameters may require an iterative algorithm. For any model, it is usually easiest to work with the logarithm of the likelihood rather than the likelihood, since likelihoods, being products of the probabilities of many data points, tend to be very small. Likelihoods multiply; log likelihoods add.

22.1 Maximum likelihood for one Gaussian We return to the Gaussian for our first examples. Assume we have data {xn }N n=1 . The log likelihood is: X √ (xn − µ)2 /(2σ 2 ). (22.1) ln P ({xn }N n=1 | µ, σ) = −N ln( 2πσ) − n

The likelihood can be expressed in terms of two functions of the data, the sample mean N X x ¯≡ xn /N, (22.2) n=1

and the sum of square deviations

S≡

X n

(xn − x ¯ )2 :

√ ln P ({xn }N ¯)2 + S]/(2σ 2 ). n=1 | µ, σ) = −N ln( 2πσ) − [N (µ − x

(22.3) (22.4)

Because the likelihood depends on the data only through x ¯ and S, these two quantities are known as sufficient statistics. Example 22.1. Differentiate the log likelihood with respect to µ and show that, if the standard deviation is known to be σ, the maximum likelihood mean µ of a Gaussian is equal to the sample mean x ¯, for any value of σ. Solution. ∂ ln P ∂µ

N (µ − x ¯) σ2 = 0 when µ = x ¯.

= −

300

(22.5) 2 (22.6)

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22.1: Maximum likelihood for one Gaussian

301

0.06 1 0.05

0.9

0.04

0.8

0.03

0.7 0.6

0.02

0.5 0.01

0.4

10

0.3 0.2

0.8 0.6 sigma

0.1 0

0.4

(a1)

0.2 0.5

0

2

1.5

1 mean

1

1.5

2

mean

0.09 sigma=0.2 sigma=0.4 sigma=0.6

4

Posterior

0.5

(a2)

4.5

mu=1 mu=1.25 mu=1.5

0.08

3.5

0.07

3

0.06

2.5

0.05

2

0.04

1.5

0.03

1

(b)

sigma

0.02

0.5

(c)

0 0

0.2

0.4

0.6

0.8

1 1.2 mean

1.4

1.6

1.8

0.01

2

0 0.2

0.4

0.6

0.8

1

1.2 1.4 1.6 1.8 2

If we Taylor-expand the log likelihood about the maximum, we can define approximate error bars on the maximum likelihood parameter: we use a quadratic approximation to estimate how far from the maximum-likelihood parameter setting we can go before the likelihood falls by some standard factor, for example e1/2 , or e4/2 . In the special case of a likelihood that is a Gaussian function of the parameters, the quadratic approximation is exact. Example 22.2. Find the second derivative of the log likelihood with respect to µ, and find the error bars on µ, given the data and σ. Solution.

∂2 N ln P = − 2 . ∂µ2 σ

2 (22.7)

Comparing this curvature with the curvature of the log of a Gaussian distribution over µ of standard deviation σ µ , exp(−µ2 /(2σµ2 )), which is −1/σµ2 , we can deduce that the error bars on µ (derived from the likelihood function) are σ σµ = √ . N

(22.8)

The error bars have this property: at the two points µ = x ¯ ± σ µ , the likelihood is smaller than its maximum value by a factor of e 1/2 . Example 22.3. Find the maximum likelihood standard deviation σ of a Gaussian, whose mean is known to be µ, in the light of data {x n }N n=1 . Find the second derivative of the log likelihood with respect to ln σ, and error bars on ln σ. Solution. The likelihood’s dependence on σ is √ Stot , (22.9) ln P ({xn }N n=1 | µ, σ) = −N ln( 2πσ) − (2σ 2 ) P where Stot = n (xn − µ)2 . To find the maximum of the likelihood, we can differentiate with respect to ln σ. [It’s often most hygienic to differentiate with

Figure 22.1. The likelihood function for the parameters of a Gaussian distribution. (a1, a2) Surface plot and contour plot of the log likelihood as a function of µ and σ. The data set of N = 5 P points had mean x ¯ = 1.0 and S = (x − x ¯)2 = 1.0. (b) The posterior probability of µ for various values of σ. (c) The posterior probability of σ for various fixed values of µ (shown as a density over ln σ).

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302

22 — Maximum Likelihood and Clustering

respect to ln u rather than u, when u is a scale variable; we use du n /d(ln u) = nun .] ∂ ln P ({xn }N Stot n=1 | µ, σ) (22.10) = −N + 2 ∂ ln σ σ This derivative is zero when Stot σ2 = , (22.11) N i.e., s PN 2 n=1 (xn − µ) σ= . (22.12) N The second derivative is Stot ∂ 2 ln P ({xn }N n=1 | µ, σ) = −2 2 , 2 ∂(ln σ) σ

(22.13)

and at the maximum-likelihood value of σ 2 , this equals −2N . So error bars on ln σ are 1 σln σ = √ . 2 (22.14) 2N . Exercise 22.4.[1 ] Show that the values of µ and lnoσ that jointly maximize the n p likelihood are: {µ, σ}ML = x ¯, σN = S/N , where σN ≡

s

PN

n=1 (xn

N

−x ¯ )2

.

(22.15)

22.2 Maximum likelihood for a mixture of Gaussians We now derive an algorithm for fitting a mixture of Gaussians to onedimensional data. In fact, this algorithm is so important to understand that, you, gentle reader, get to derive the algorithm. Please work through the following exercise. Exercise 22.5. [2, p.310] A random variable x is assumed to have a probability distribution that is a mixture of two Gaussians, " 2 #  X (x − µk )2 1 P (x | µ1 , µ2 , σ) = , (22.16) pk √ exp − 2σ 2 2πσ 2 k=1

where the two Gaussians are given the labels k = 1 and k = 2; the prior probability of the class label k is {p 1 = 1/2, p2 = 1/2}; {µk } are the means of the two Gaussians; and both have standard deviation σ. For brevity, we denote these parameters by θ ≡ {{µk }, σ}. A data set consists of N points {xn }N n=1 which are assumed to be independent samples from this distribution. Let k n denote the unknown class label of the nth point. Assuming that {µk } and σ are known, show that the posterior probability of the class label kn of the nth point can be written as P (kn = 1 | xn , θ) =

1 1 + exp[−(w1 xn + w0 )]

P (kn = 2 | xn , θ) =

1 , 1 + exp[+(w1 xn + w0 )]

(22.17)

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22.3: Enhancements to soft K-means

303

and give expressions for w1 and w0 . Assume now that the means {µk } are not known, and that we wish to infer them from the data {xn }N n=1 . (The standard deviation σ is known.) In the remainder of this question we will derive an iterative algorithm for finding values for {µk } that maximize the likelihood, Y P ({xn }N P (xn | {µk }, σ). (22.18) n=1 | {µk }, σ) = n

Let L denote the natural log of the likelihood. Show that the derivative of the log likelihood with respect to µk is given by X (xn − µk ) ∂ pk|n L= , ∂µk σ2 n

(22.19)

where pk|n ≡ P (kn = k | xn , θ) appeared above at equation (22.17). Show, neglecting terms in ∂µ∂ k P (kn = k | xn , θ), that the second derivative is approximately given by X ∂2 1 pk|n 2 . L = − σ ∂µ2k n

(22.20)

Hence show that from an initial state µ 1 , µ2 , an approximate Newton–Raphson step updates these parameters to µ01 , µ02 , where P n pk|n xn 0 . (22.21) µk = P n pk|n

[The method for maximizing L(µ) updates µ to µ 0 = µ − h .Newton–Raphson i 2 ∂L ∂ L .] ∂µ ∂µ2

0

1

2

3

4

5

6

Assuming that σ = 1, sketch a contour plot of the likelihood function as a function of µ1 and µ2 for the data set shown above. The data set consists of 32 points. Describe the peaks in your sketch and indicate their widths. Notice that the algorithm you have derived for maximizing the likelihood is identical to the soft K-means algorithm of section 20.4. Now that it is clear that clustering can be viewed as mixture-density-modelling, we are able to derive enhancements to the K-means algorithm, which rectify the problems we noted earlier.

22.3 Enhancements to soft K-means Algorithm 22.2 shows a version of the soft-K-means algorithm corresponding to a modelling assumption that each cluster is a spherical Gaussian having its own width (each cluster has its own β (k) = 1/ σk2 ). The algorithm updates the lengthscales σk for itself. The algorithm also includes cluster weight parameters π1 , π2 , . . . , πK which also update themselves, allowing accurate modelling of data from clusters of unequal weights. This algorithm is demonstrated in figure 22.3 for two data sets that we’ve seen before. The second example shows

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304

22 — Maximum Likelihood and Clustering Algorithm 22.2. The soft K-means algorithm, version 2.

Assignment step. The responsibilities are   1 (k) (n) 1 d(m , x ) exp − πk (√2πσ I k) σ2 (n)   k rk = P 1 (k 0 ) (n) √ 1 π d(m , x ) exp − 0 k k ( 2πσk0 )I σk20

(22.22)

where I is the dimensionality of x.

Update step. Each cluster’s parameters, m (k) , πk , and σk2 , are adjusted to match the data points that it is responsible for. X (n) rk x(n) n

m(k) =

σk2 =

X n

(22.23)

R(k)

(n)

rk (x(n) − m(k) )2

(22.24)

IR(k)

R(k) πk = P (k) kR

(22.25)

where R(k) is the total responsibility of mean k, X (n) R(k) = rk .

(22.26)

n

(n)

rk

t=0

t=1

t=2

t=3

t=9

t=0

t=1

t = 10

t = 20

t = 30

=

πk Q

I . X 1 (k) (n) 2 (k) exp − (m − x ) 2(σi )2 √ i i (k) I 2πσ i=1 i=1 i P (numerator, with k 0 in place of k) 0 k

σi2

(k)

=

X n

(n)

(n)

rk (xi

t = 35

Algorithm 22.4. The soft K-means algorithm, version 3, which corresponds to a model of axis-aligned Gaussians.

(22.27)

(k)

− m i )2

R(k)

!

Figure 22.3. Soft K-means algorithm, with K = 2, applied (a) to the 40-point data set of figure 20.3; (b) to the little ’n’ large data set of figure 20.5.

(22.28)

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22.4: A fatal flaw of maximum likelihood t = 10

t=0

t=0

t = 10

305 t = 20

t = 20

t = 26

t = 30

Figure 22.5. Soft K-means algorithm, version 3, applied to the data consisting of two cigar-shaped clusters. K = 2 (cf. figure 20.6).

t = 32

Figure 22.6. Soft K-means algorithm, version 3, applied to the little ’n’ large data set. K = 2.

that convergence can take a long time, but eventually the algorithm identifies the small cluster and the large cluster. Soft K-means, version 2, is a maximum-likelihood algorithm for fitting a mixture of spherical Gaussians to data – ‘spherical’ meaning that the variance of the Gaussian is the same in all directions. This algorithm is still no good at modelling the cigar-shaped clusters of figure 20.6. If we wish to model the clusters by axis-aligned Gaussians with possibly-unequal variances, we replace the assignment rule (22.22) and the variance update rule (22.24) by the rules (22.27) and (22.28) displayed in algorithm 22.4. This third version of soft K-means is demonstrated in figure 22.5 on the ‘two cigars’ data set of figure 20.6. After 30 iterations, the algorithm correctly locates the two clusters. Figure 22.6 shows the same algorithm applied to the little ’n’ large data set; again, the correct cluster locations are found.

A proof that the algorithm does indeed maximize the likelihood is deferred to section 33.7.

22.4 A fatal flaw of maximum likelihood Finally, figure 22.7 sounds a cautionary note: when we fit K = 4 means to our first toy data set, we sometimes find that very small clusters form, covering just one or two data points. This is a pathological property of soft K-means clustering, versions 2 and 3. . Exercise 22.6.[2 ] Investigate what happens if one mean m (k) sits exactly on top of one data point; show that if the variance σ k2 is sufficiently small, then no return is possible: σk2 becomes ever smaller. t=0

t=5

t = 10

t = 20

Figure 22.7. Soft K-means algorithm applied to a data set of 40 points. K = 4. Notice that at convergence, one very small cluster has formed between two data points.

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306

22 — Maximum Likelihood and Clustering

KABOOM! Soft K-means can blow up. Put one cluster exactly on one data point and let its variance go to zero – you can obtain an arbitrarily large likelihood! Maximum likelihood methods can break down by finding highly tuned models that fit part of the data perfectly. This phenomenon is known as overfitting. The reason we are not interested in these solutions with enormous likelihood is this: sure, these parameter-settings may have enormous posterior probability density, but the density is large over only a very small volume of parameter space. So the probability mass associated with these likelihood spikes is usually tiny. We conclude that maximum likelihood methods are not a satisfactory general solution to data-modelling problems: the likelihood may be infinitely large at certain parameter settings. Even if the likelihood does not have infinitelylarge spikes, the maximum of the likelihood is often unrepresentative, in highdimensional problems. Even in low-dimensional problems, maximum likelihood solutions can be unrepresentative. As you may know from basic statistics, the maximum likelihood estimator (22.15) for a Gaussian’s standard deviation, σ N , is a biased estimator, a topic that we’ll take up in Chapter 24. The maximum a posteriori (MAP) method A popular replacement for maximizing the likelihood is maximizing the Bayesian posterior probability density of the parameters instead. However, multiplying the likelihood by a prior and maximizing the posterior does not make the above problems go away; the posterior density often also has infinitely-large spikes, and the maximum of the posterior probability density is often unrepresentative of the whole posterior distribution. Think back to the concept of typicality, which we encountered in Chapter 4: in high dimensions, most of the probability mass is in a typical set whose properties are quite different from the points that have the maximum probability density. Maxima are atypical. A further reason for disliking the maximum a posteriori is that it is basisdependent. If we make a nonlinear change of basis from the parameter θ to the parameter u = f (θ) then the probability density of θ is transformed to ∂θ (22.29) P (u) = P (θ) . ∂u The maximum of the density P (u) will usually not coincide with the maximum of the density P (θ). (For figures illustrating such nonlinear changes of basis, see the next chapter.) It seems undesirable to use a method whose answers change when we change representation.

Further reading The soft K-means algorithm is at the heart of the automatic classification package, AutoClass (Hanson et al., 1991b; Hanson et al., 1991a).

22.5 Further exercises Exercises where maximum likelihood may be useful Exercise 22.7.[3 ] Make a version of the K-means algorithm that models the data as a mixture of K arbitrary Gaussians, i.e., Gaussians that are not constrained to be axis-aligned.

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22.5: Further exercises

307

. Exercise 22.8.[2 ] (a) A photon counter is pointed at a remote star for one minute, in order to infer the brightness, i.e., the rate of photons arriving at the counter per minute, λ. Assuming the number of photons collected r has a Poisson distribution with mean λ, P (r | λ) = exp(−λ)

λr , r!

(22.30)

what is the maximum likelihood estimate for λ, given r = 9? Find error bars on ln λ. (b) Same situation, but now we assume that the counter detects not only photons from the star but also ‘background’ photons. The background rate of photons is known to be b = 13 photons per minute. We assume the number of photons collected, r, has a Poisson distribution with mean λ+b. Now, given r = 9 detected photons, what is the maximum likelihood estimate for λ? Comment on this answer, discussing also the Bayesian posterior distribution, and the ˆ ≡ r − b. ‘unbiased estimator’ of sampling theory, λ Exercise 22.9.[2 ] A bent coin is tossed N times, giving N a heads and Nb tails. Assume a beta distribution prior for the probability of heads, p, for example the uniform distribution. Find the maximum likelihood and maximum a posteriori values of p, then find the maximum likelihood and maximum a posteriori values of the logit a ≡ ln[p/(1−p)]. Compare with the predictive distribution, i.e., the probability that the next toss will come up heads. . Exercise 22.10. [2 ] Two men looked through prison bars; one saw stars, the other tried to infer where the window frame was. From the other side of a room, you look through a window and see stars at locations {(xn , yn )}. You can’t see the window edges because it is totally dark apart from the stars. Assuming the window is rectangular and that the visible stars’ locations are independently randomly distributed, what are the inferred values of (xmin , ymin , xmax , ymax ), according to maximum likelihood? Sketch the likelihood as a function of x max , for fixed xmin , ymin , and ymax . . Exercise 22.11. [3 ] A sailor infers his location (x, y) by measuring the bearings of three buoys whose locations (xn , yn ) are given on his chart. Let the true bearings of the buoys be θn . Assuming that his measurement θ˜n of each bearing is subject to Gaussian noise of small standard deviation σ, what is his inferred location, by maximum likelihood? The sailor’s rule of thumb says that the boat’s position can be taken to be the centre of the cocked hat, the triangle produced by the intersection of the three measured bearings (figure 22.8). Can you persuade him that the maximum likelihood answer is better? . Exercise 22.12. [3, p.310] Maximum likelihood fitting of an exponential-family model. Assume that a variable x comes from a probability distribution of the form ! X 1 exp wk fk (x) , (22.31) P (x | w) = Z(w) k

(xmax , ymax )

? ?

? ??

?

(xmin , ymin )

b

(x3 , y3 )Q

(x1 , y1 )

b

QA Q AQ A Q A Q A A A A Ab (x2 , y2 )

Figure 22.8. The standard way of drawing three slightly inconsistent bearings on a chart produces a triangle called a cocked hat. Where is the sailor?

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308

22 — Maximum Likelihood and Clustering where the functions fk (x) are given, and the parameters w = {w k } are not known. A data set {x(n) } of N points is supplied. Show by differentiating the log likelihood that the maximum-likelihood parameters wML satisfy X x

P (x | wML )fk (x) =

1 X fk (x(n) ), N n

(22.32)

where the left-hand sum is over all x, and the right-hand sum is over the data points. A shorthand for this result is that each function-average under the fitted model must equal the function-average found in the data: hfk iP (x | wML ) = hfk iData .

(22.33)

. Exercise 22.13. [3 ] ‘Maximum entropy’ fitting of models to constraints. When confronted by a probability distribution P (x) about which only a few facts are known, the maximum entropy principle (maxent) offers a rule for choosing a distribution that satisfies those constraints. According to maxent, you should select the P (x) that maximizes the entropy H=

X

P (x) log 1/P (x),

(22.34)

x

subject to the constraints. Assuming the constraints assert that the averages of certain functions fk (x) are known, i.e., hfk iP (x) = Fk ,

(22.35)

show, by introducing Lagrange multipliers (one for each constraint, including normalization), that the maximum-entropy distribution has the form ! X 1 wk fk (x) , (22.36) P (x)Maxent = exp Z k

where the parameters Z and {wk } are set such that the constraints (22.35) are satisfied. And hence the maximum entropy method gives identical results to maximum likelihood fitting of an exponential-family model (previous exercise). The maximum entropy method has sometimes been recommended as a method for assigning prior distributions in Bayesian modelling. While the outcomes of the maximum entropy method are sometimes interesting and thought-provoking, I do not advocate maxent as the approach to assigning priors. Maximum entropy is also sometimes proposed as a method for solving inference problems – for example, ‘given that the mean score of this unfair six-sided die is 2.5, what is its probability distribution (p1 , p2 , p3 , p4 , p5 , p6 )?’ I think it is a bad idea to use maximum entropy in this way; it can give silly answers. The correct way to solve inference problems is to use Bayes’ theorem.

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22.5: Further exercises

309

Exercises where maximum likelihood and MAP have difficulties . Exercise 22.14. [2 ] This exercise explores the idea that maximizing a probability density is a poor way to find a point that is representative of the density. √ Consider a Gaussian in a k-dimensional space, P distribution 2 ). Show that nearly all of the P (w) = (1/ 2π σW )k exp(− k1 wi2 /2σW √ probability mass of a Gaussian is in√a thin shell of radius r = kσW and of thickness proportional to r/ k. For example, in 1000 dimensions, 90% of the mass of a Gaussian with σ W = 1 is in a shell of radius 31.6 and thickness 2.8. However, the probability density at the origin is ek/2 ' 10217 times bigger than the density at this shell where most of the probability mass is. Now consider two Gaussian densities in 1000 dimensions that differ in radius σW by just 1%, and that contain equal total probability mass. Show that the maximum probability density is greater at the centre of the Gaussian with smaller σW by a factor of ∼ exp(0.01k) ' 20 000.

In ill-posed problems, a typical posterior distribution is often a weighted superposition of Gaussians with varying means and standard deviations, so the true posterior has a skew peak, with the maximum of the probability density located near the mean of the Gaussian distribution that has the smallest standard deviation, not the Gaussian with the greatest weight. . Exercise 22.15. [3 ] The seven scientists. N datapoints {x n } are drawn from N distributions, all of which are Gaussian with a common mean µ but with different unknown standard deviations σ n . What are the maximum likelihood parameters µ, {σn } given the data? For example, seven scientists (A, B, C, D, E, F, G) with wildly-differing experimental skills measure µ. You expect some of them to do accurate work (i.e., to have small σn ), and some of them to turn in wildly inaccurate answers (i.e., to have enormous σn ). Figure 22.9 shows their seven results. What is µ, and how reliable is each scientist? I hope you agree that, intuitively, it looks pretty certain that A and B are both inept measurers, that D–G are better, and that the true value of µ is somewhere close to 10. But what does maximizing the likelihood tell you? Exercise 22.16. [3 ] Problems with MAP method. A collection of widgets i = 1, . . . , k have a property called ‘wodge’, w i , which we measure, widget by widget, in noisy experiments with a known noise level σ ν = 1.0. Our model for these quantities is that they come from a Gaussian prior 2 is not known. Our prior for P (wi | α) = Normal(0, 1/α), where α = 1/σW this variance is flat over log σW from σW = 0.1 to σW = 10. Scenario 1. Suppose four widgets have been measured and give the following data: {d1 , d2 , d3 , d4 } = {2.2, −2.2, 2.8, −2.8}. We are interested in inferring the wodges of these four widgets. (a) Find the values of w and α that maximize the posterior probability P (w, log α | d).

(b) Marginalize over α and find the posterior probability density of w given the data. [Integration skills required. See MacKay (1999a) for solution.] Find maxima of P (w | d). [Answer: two maxima – one at wMP = {1.8, −1.8, 2.2, −2.2}, with error bars on all four parameters

A

-30

B C D-G

-20

-10

0

10

Scientist

xn

A B C D E F G

−27.020 3.570 8.191 9.898 9.603 9.945 10.056

20

Figure 22.9. Seven measurements {xn } of a parameter µ by seven scientists each having his own noise-level σn .

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310

22 — Maximum Likelihood and Clustering (obtained from Gaussian approximation to the posterior) ±0.9; and 0 one at wMP = {0.03, −0.03, 0.04, −0.04} with error bars ±0.1.] Scenario 2. Suppose in addition to the four measurements above we are now informed that there are four more widgets that have been measured with a much less accurate instrument, having σ ν0 = 100.0. Thus we now have both well-determined and ill-determined parameters, as in a typical ill-posed problem. The data from these measurements were a string of uninformative values, {d5 , d6 , d7 , d8 } = {100, −100, 100, −100}.

We are again asked to infer the wodges of the widgets. Intuitively, our inferences about the well-measured widgets should be negligibly affected by this vacuous information about the poorly-measured widgets. But what happens to the MAP method? (a) Find the values of w and α that maximize the posterior probability P (w, log α | d).

5 4

(b) Find maxima of P (w | d). [Answer: only one maximum, w MP = {0.03, −0.03, 0.03, −0.03, 0.0001, −0.0001, 0.0001, −0.0001}, with error bars on all eight parameters ±0.11.]

3 2 1

22.6 Solutions 0

Solution to exercise 22.5 (p.302). Figure 22.10 shows a contour plot of the likelihood function for the 32 data points. The peaks are pretty-near centred on the points (1, 5) and (5, 1), and are pretty-near circular in√their contours. The width of each of the peaks is a standard deviation of σ/ 16 = 1/4. The peaks are roughly Gaussian in shape. Solution to exercise 22.12 (p.307).

The log likelihood is: XX ln P ({x(n) } | w) = −N ln Z(w) + wk fk (x(n) ). n

(22.37)

k

X ∂ ∂ fk (x). ln P ({x(n) } | w) = −N ln Z(w) + ∂wk ∂wk n

(22.38)

Now, the fun part is what happens when we differentiate the log of the normalizing constant: ! X 1 X ∂ ∂ ln Z(w) = exp wk0 fk0 (x) ∂wk Z(w) x ∂wk 0 k

=

X X 1 X exp wk0 fk0 (x) fk (x) = P (x | w)fk (x), Z(w) x 0 x k

so

!

X X ∂ ln P ({x(n) } | w) = −N P (x | w)fk (x) + fk (x), ∂wk x n

(22.39)

(22.40)

and at the maximum of the likelihood, X x

P (x | wML )fk (x) =

1 X fk (x(n) ). N n

(22.41)

1

2

3

4

5

0

Figure 22.10. The likelihood as a function of µ1 and µ2 .

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23 Useful Probability Distributions In Bayesian data modelling, there’s a small collection of probability distributions that come up again and again. The purpose of this chapter is to introduce these distributions so that they won’t be intimidating when encountered in combat situations. There is no need to memorize any of them, except perhaps the Gaussian; if a distribution is important enough, it will memorize itself, and otherwise, it can easily be looked up.

0.3 0.25 0.2 0.15 0.1 0.05 0 0 1 2 3 4 5 6 7 8 9 10

1 0.1 0.01 0.001

23.1 Distributions over integers

0.0001 1e-05

Binomial, Poisson, exponential We already encountered the binomial distribution and the Poisson distribution on page 2. The binomial distribution for an integer r with parameters f (the bias, f ∈ [0, 1]) and N (the number of trials) is:   N r P (r | f, N ) = f (1 − f )N −r r ∈ {0, 1, 2, . . . , N }. (23.1) r The binomial distribution arises, for example, when we flip a bent coin, with bias f , N times, and observe the number of heads, r.

λr r!

311

0.25 0.2 0.15

r ∈ (0, 1, 2, . . . , ∞),

r ∈ (0, 1, 2, . . . , ∞),

0

(23.2)

(23.3)

(23.4)

5

10

15

5

10

15

1 0.1 0.01 0.001 0.0001 1e-05 1e-06 1e-07 0

arises in waiting problems. How long will you have to wait until a six is rolled, if a fair six-sided dice is rolled? Answer: the probability distribution of the number of rolls, r, is exponential over integers with parameter f = 5/6. The distribution may also be written

where λ = ln(1/f ).

Figure 23.1. The binomial distribution P (r | f = 0.3, N = 10), on a linear scale (top) and a logarithmic scale (bottom).

0

r ∈ {0, 1, 2, . . .}.

The exponential distribution on integers,,

P (r | f ) = (1 − f ) e−λr

r

0.05

The Poisson distribution arises, for example, when we count the number of photons r that arrive in a pixel during a fixed interval, given that the mean intensity on the pixel corresponds to an average number of photons λ.

P (r | f ) = f r (1 − f )

0 1 2 3 4 5 6 7 8 9 10

0.1

The Poisson distribution with parameter λ > 0 is: P (r | λ) = e−λ

1e-06

r

Figure 23.2. The Poisson distribution P (r | λ = 2.7), on a linear scale (top) and a logarithmic scale (bottom).

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312

23 — Useful Probability Distributions

23.2 Distributions over unbounded real numbers Gaussian, Student, Cauchy, biexponential, inverse-cosh. The Gaussian distribution or normal distribution with mean µ and standard deviation σ is   1 (x − µ)2 x ∈ (−∞, ∞), (23.5) P (x | µ, σ) = exp − Z 2σ 2 where Z=

√ 2πσ 2 .

(23.6)

It is sometimes useful to work with the quantity τ ≡ 1/σ 2 , which is called the precision parameter of the Gaussian. A sample z from a standard univariate Gaussian can be generated by computing p z = cos(2πu1 ) 2 ln(1/u2 ), (23.7)

where u1 p and u2 are uniformly distributed in (0, 1). A second sample z 2 = sin(2πu1 ) 2 ln(1/u2 ), independent of the first, can then be obtained for free. The Gaussian distribution is widely used and often asserted to be a very common distribution in the real world, but I am sceptical about this assertion. Yes, unimodal distributions may be common; but a Gaussian is a special, rather extreme, unimodal distribution. It has very light tails: the logprobability-density decreases quadratically. The typical deviation of x from µ is σ, but the respective probabilities that x deviates from µ by more than 2σ, 3σ, 4σ, and 5σ, are 0.046, 0.003, 6 × 10 −5 , and 6 × 10−7 . In my experience, deviations from a mean four or five times greater than the typical deviation may be rare, but not as rare as 6 × 10−5 ! I therefore urge caution in the use of Gaussian distributions: if a variable that is modelled with a Gaussian actually has a heavier-tailed distribution, the rest of the model will contort itself to reduce the deviations of the outliers, like a sheet of paper being crushed by a rubber band. [1 ]

. Exercise 23.1. Pick a variable that is supposedly bell-shaped in probability distribution, gather data, and make a plot of the variable’s empirical distribution. Show the distribution as a histogram on a log scale and investigate whether the tails are well-modelled by a Gaussian distribution. [One example of a variable to study is the amplitude of an audio signal.] One distribution with heavier tails than a Gaussian is a mixture of Gaussians. A mixture of two Gaussians, for example, is defined by two means, two standard deviations, and two mixing coefficients π 1 and π2 , satisfying π1 + π2 = 1, πi ≥ 0.     2 π1 π2 (x−µ2 )2 1) √ P (x | µ1 , σ1 , π1 , µ2 , σ2 , π2 ) = √ exp − (x−µ + exp − . 2 2 2σ1 2σ2 2πσ1 2πσ2 If we take an appropriately weighted mixture of an infinite number of Gaussians, all having mean µ, we obtain a Student-t distribution, P (x | µ, s, n) = where Z=

1 1 , Z (1 + (x − µ)2 /(ns2 ))(n+1)/2

(23.8)

√ πns2

(23.9)

Γ(n/2) Γ((n + 1)/2)

0.5 0.4 0.3 0.2 0.1 0 -2

0

2

4

6

8

2

4

6

8

0.1 0.01 0.001 0.0001 -2

0

Figure 23.3. Three unimodal distributions. Two Student distributions, with parameters (m, s) = (1, 1) (heavy line) (a Cauchy distribution) and (2, 4) (light line), and a Gaussian distribution with mean µ = 3 and standard deviation σ = 3 (dashed line), shown on linear vertical scales (top) and logarithmic vertical scales (bottom). Notice that the heavy tails of the Cauchy distribution are scarcely evident in the upper ‘bell-shaped curve’.

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23.3: Distributions over positive real numbers

313

and n is called the number of degrees of freedom and Γ is the gamma function. If n > 1 then the Student distribution (23.8) has a mean and that mean is µ. If n > 2 the distribution also has a finite variance, σ 2 = ns2 /(n − 2). As n → ∞, the Student distribution approaches the normal distribution with mean µ and standard deviation s. The Student distribution arises both in classical statistics (as the sampling-theoretic distribution of certain statistics) and in Bayesian inference (as the probability distribution of a variable coming from a Gaussian distribution whose standard deviation we aren’t sure of). In the special case n = 1, the Student distribution is called the Cauchy distribution. A distribution whose tails are intermediate in heaviness between Student and Gaussian is the biexponential distribution,   |x − µ| 1 x ∈ (−∞, ∞) (23.10) P (x | µ, s) = exp − Z s where Z = 2s.

(23.11)

The inverse-cosh distribution P (x | β) ∝

1 [cosh(βx)]1/β

(23.12)

is a popular model in independent component analysis. In the limit of large β, the probability distribution P (x | β) becomes a biexponential distribution. In the limit β → 0 P (x | β) approaches a Gaussian with mean zero and variance 1/β.

23.3 Distributions over positive real numbers Exponential, gamma, inverse-gamma, and log-normal. The exponential distribution,  x 1 P (x | s) = exp − x ∈ (0, ∞), Z s

(23.13)

where

Z = s,

(23.14)

arises in waiting problems. How long will you have to wait for a bus in Poissonville, given that buses arrive independently at random with one every s minutes on average? Answer: the probability distribution of your wait, x, is exponential with mean s. The gamma distribution is like a Gaussian distribution, except whereas the Gaussian goes from −∞ to ∞, gamma distributions go from 0 to ∞. Just as the Gaussian distribution has two parameters µ and σ which control the mean and width of the distribution, the gamma distribution has two parameters. It is the product of the one-parameter exponential distribution (23.13) with a polynomial, xc−1 . The exponent c in the polynomial is the second parameter.  x 1  x c−1 P (x | s, c) = Γ(x; s, c) = , 0≤x<∞ (23.15) exp − Z s s where

Z = Γ(c)s.

(23.16)

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314

23 — Useful Probability Distributions 1 0.9 0.8 0.7 0.6 0.5 0.4 0.3 0.2 0.1 0

Figure 23.4. Two gamma distributions, with parameters (s, c) = (1, 3) (heavy lines) and 10, 0.3 (light lines), shown on linear vertical scales (top) and logarithmic vertical scales (bottom); and shown as a function of x on the left (23.15) and l = ln x on the right (23.18).

0.8 0.7 0.6 0.5 0.4 0.3 0.2 0.1 0 0

2

4

6

8

10

-4

-2

0

2

4

2

4

1 0.1

0.1

0.01

0.01

0.001

0.001

0.0001

0.0001 0

2

4

6

8

10

x

-4

-2

0

l = ln x

This is a simple peaked distribution with mean sc and variance s 2 c. It is often natural to represent a positive real variable x in terms of its logarithm l = ln x. The probability density of l is ∂x (23.17) P (l) = P (x(l)) = P (x(l))x(l) ∂l   c  1 x(l) x(l) = exp − , (23.18) Zl s s where Zl = Γ(c).

(23.19)

[The gamma distribution is named after its normalizing constant – an odd convention, it seems to me!] Figure 23.4 shows a couple of gamma distributions as a function of x and of l. Notice that where the original gamma distribution (23.15) may have a ‘spike’ at x = 0, the distribution over l never has such a spike. The spike is an artefact of a bad choice of basis. In the limit sc = 1, c → 0, we obtain the noninformative prior for a scale parameter, the 1/x prior. This improper prior is called noninformative because it has no associated length scale, no characteristic value of x, so it prefers all values of x equally. It is invariant under the reparameterization x = mx. If we transform the 1/x probability density into a density over l = ln x we find the latter density is uniform. . Exercise 23.2.[1 ] Imagine that we reparameterize a positive variable x in terms of its cube root, u = x1/3 . If the probability density of x is the improper distribution 1/x, what is the probability density of u? The gamma distribution is always a unimodal density over l = ln x, and, as can be seen in the figures, it is asymmetric. If x has a gamma distribution, and we decide to work in terms of the inverse of x, v = 1/x, we obtain a new distribution, in which the density over l is flipped left-for-right: the probability density of v is called an inverse-gamma distribution, P (v | s, c) =

1 Zv



1 sv

c+1

  1 exp − , sv

0≤v<∞

(23.20)

where Zv = Γ(c)/s.

(23.21)

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23.4: Distributions over periodic variables

315 Figure 23.5. Two inverse gamma distributions, with parameters (s, c) = (1, 3) (heavy lines) and 10, 0.3 (light lines), shown on linear vertical scales (top) and logarithmic vertical scales (bottom); and shown as a function of x on the left and l = ln x on the right.

0.8 0.7 0.6 0.5 0.4 0.3 0.2 0.1 0

2.5 2 1.5 1 0.5 0 0

1

2

3

-4

-2

0

2

0

2

4

1 0.1

0.1

0.01

0.01 0.001

0.001

0.0001

0.0001 0

1

2

3

-4

v

-2

4

ln v

Gamma and inverse gamma distributions crop up in many inference problems in which a positive quantity is inferred from data. Examples include inferring the variance of Gaussian noise from some noise samples, and inferring the rate parameter of a Poisson distribution from the count. Gamma distributions also arise naturally in the distributions of waiting times between Poisson-distributed events. Given a Poisson process with rate λ, the probability density of the arrival time x of the mth event is λ(λx)m−1 −λx e . (m−1)!

(23.22)

Log-normal distribution Another distribution over a positive real number x is the log-normal distribution, which is the distribution that results when l = ln x has a normal distribution. We define m to be the median value of x, and s to be the standard deviation of ln x.   1 (l − ln m)2 l ∈ (−∞, ∞), (23.23) P (l | m, s) = exp − Z 2s2 where

√ Z = 2πs2 ,

(23.24)

implies P (x | m, s) =

  11 (ln x − ln m)2 exp − xZ 2s2

x ∈ (0, ∞).

(23.25)

23.4 Distributions over periodic variables A periodic variable θ is a real number ∈ [0, 2π] having the property that θ = 0 and θ = 2π are equivalent. A distribution that plays for periodic variables the role played by the Gaussian distribution for real variables is the Von Mises distribution: P (θ | µ, β) =

1 exp (β cos(θ − µ)) Z

θ ∈ (0, 2π).

(23.26)

The normalizing constant is Z = 2πI0 (β), where I0 (x) is a modified Bessel function.

0.4 0.35 0.3 0.25 0.2 0.15 0.1 0.05 0 0

1

2

3

4

5

0

1

2

3

4

5

0.1 0.01 0.001 0.0001

Figure 23.6. Two log-normal distributions, with parameters (m, s) = (3, 1.8) (heavy line) and (3, 0.7) (light line), shown on linear vertical scales (top) and logarithmic vertical scales (bottom). [Yes, they really do have the same value of the median, m = 3.]

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316

23 — Useful Probability Distributions

A distribution that arises from Brownian diffusion around the circle is the wrapped Gaussian distribution, P (θ | µ, σ) =

∞ X

n=−∞

5 4 3

2

Normal(θ; (µ + 2πn), σ ) θ ∈ (0, 2π).

(23.27)

2 1 0 0

23.5 Distributions over probabilities

0.25

0.5

0.75

1

0.6 0.5

Beta distribution, Dirichlet distribution, entropic distribution The beta distribution is a probability density over a variable p that is a probability, p ∈ (0, 1): 1 pu1 −1 (1 − p)u2 −1 . P (p | u1 , u2 ) = Z(u1 , u2 )

(23.28)

The parameters u1 , u2 may take any positive value. The normalizing constant is the beta function, Γ(u1 )Γ(u2 ) Z(u1 , u2 ) = . (23.29) Γ(u1 + u2 ) Special cases include the uniform distribution – u 1 = 1, u2 = 1; the Jeffreys prior – u1 = 0.5, u2 = 0.5; and the improper Laplace prior – u 1 = 0, u2 = 0. If we transform the beta distribution to the corresponding density over the logit l ≡ ln p/ (1 − p), we find it is always a pleasant bell-shaped density over l, while the density over p may have singularities at p = 0 and p = 1 (figure 23.7).

0.3 0.2 0.1 0 -6

-4

-2

0

2

4

6

Figure 23.7. Three beta distributions, with (u1 , u2 ) = (0.3, 1), (1.3, 1), and (12, 2). The upper figure shows P (p | u1 , u2 ) as a function of p; the lower shows the corresponding density over the logit, ln

p . 1−p

Notice how well-behaved the densities are as a function of the logit.

More dimensions The Dirichlet distribution is a density over an I-dimensional vector p whose I components are positive and sum to 1. The beta distribution is a special case of a Dirichlet distribution with I = 2. The Dirichlet distribution is parameterized by a measure u (a vector with all coefficients u i > 0) which I will write here P as u = αm, where m is a normalized measure over the I components ( mi = 1), and α is positive: I

P (p | αm) =

0.4

Y αm −1 P 1 pi i δ ( i pi − 1) ≡ Dirichlet (I) (p | αm). (23.30) Z(αm) i=1

The function δ(x) is the Dirac delta function, which P restricts the distribution to the simplex such that p is normalized, i.e., i pi = 1. The normalizing constant of the Dirichlet distribution is: Y Z(αm) = Γ(αmi ) /Γ(α) . (23.31) i

The vector m is the mean of the probability distribution: Z Dirichlet (I) (p | αm) p dIp = m.

(23.32)

When working with a probability vector p, it is often helpful to work in the ‘softmax basis’, in which, for example, a three-dimensional probability p = (p1 , p2 , p3 ) is represented by three numbers a1 , a2 , a3 satisfying a1 +a2 +a3 = 0 and P 1 (23.33) pi = eai , where Z = i eai . Z This nonlinear transformation is analogous to the σ → ln σ transformation for a scale variable and the logit transformation for a single probability, p →

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23.5: Distributions over probabilities u = (20, 10, 7)

317

u = (0.2, 1, 2)

u = (0.2, 0.3, 0.15)

8

8

8

4

4

4

0

0

0

-4

-4

-4

-8

-8 -8

-4

0

4

8

Figure 23.8. Three Dirichlet distributions over a three-dimensional probability vector (p1 , p2 , p3 ). The upper figures show 1000 random draws from each distribution, showing the values of p1 and p2 on the two axes. p3 = 1 − (p1 + p2 ). The triangle in the first figure is the simplex of legal probability distributions. The lower figures show the same points in the ‘softmax’ basis (equation (23.33)). The two axes show a1 and a2 . a3 = −a1 − a2 .

-8 -8

-4

0

4

8

-8

-4

0

4

8

p . In the softmax basis, the ugly minus-ones in the exponents in the ln 1−p Dirichlet distribution (23.30) disappear, and the density is given by: I

P (a | αm) ∝

Y αm P 1 pi i δ ( i ai ) . Z(αm)

(23.34)

i=1

The role of the parameter α can be characterized in two ways. First, α measures the sharpness of the distribution (figure 23.8); it measures how different we expect typical samples p from the distribution to be from the mean m, just as the precision τ = 1/σ 2 of a Gaussian measures how far samples stray from its mean. A large value of α produces a distribution over p that is sharply peaked around m. The effect of α in higher-dimensional situations can be visualized by drawing a typical sample from the distribution Dirichlet (I) (p | αm), with m set to the uniform vector mi = 1/I , and making a Zipf plot, that is, a ranked plot of the values of the components p i . It is traditional to plot both pi (vertical axis) and the rank (horizontal axis) on logarithmic scales so that power law relationships appear as straight lines. Figure 23.9 shows these plots for a single sample from ensembles with I = 100 and I = 1000 and with α from 0.1 to 1000. For large α, the plot is shallow with many components having similar values. For small α, typically one component p i receives an overwhelming share of the probability, and of the small probability that remains to be shared among the other components, another component p i0 receives a similarly large share. In the limit as α goes to zero, the plot tends to an increasingly steep power law. Second, we can characterize the role of α in terms of the predictive distribution that results when we observe samples from p and obtain counts F = (F1 , F2 , . . . , FI ) of the possible outcomes. The value of α defines the number of samples from p that are required in order that the data dominate over the prior in predictions. Exercise 23.3.[3 ] The Dirichlet distribution satisfies a nice additivity property. Imagine that a biased six-sided die has two red faces and four blue faces. The die is rolled N times and two Bayesians examine the outcomes in order to infer the bias of the die and make predictions. One Bayesian has access to the red/blue colour outcomes only, and he infers a twocomponent probability vector (pR , pB ). The other Bayesian has access to each full outcome: he can see which of the six faces came up, and he infers a six-component probability vector (p 1 , p2 , p3 , p4 , p5 , p6 ), where

I = 100 1 0.1 1 10 100 1000

0.1

0.01

0.001

0.0001 1

10

100

I = 1000 1 0.1 1 10 100 1000

0.1

0.01

0.001

0.0001

1e-05 1

10

100

1000

Figure 23.9. Zipf plots for random samples from Dirichlet distributions with various values of α = 0.1 . . . 1000. For each value of I = 100 or 1000 and each α, one sample p from the Dirichlet distribution was generated. The Zipf plot shows the probabilities pi , ranked by magnitude, versus their rank.

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318

23 — Useful Probability Distributions pR = p1 + p2 and pB = p3 + p4 + p5 + p6 . Assuming that the second Bayesian assigns a Dirichlet distribution to (p 1 , p2 , p3 , p4 , p5 , p6 ) with hyperparameters (u1 , u2 , u3 , u4 , u5 , u6 ), show that, in order for the first Bayesian’s inferences to be consistent with those of the second Bayesian, the first Bayesian’s prior should be a Dirichlet distribution with hyperparameters ((u1 + u2 ), (u3 + u4 + u5 + u6 )). Hint: R 6 a brute-force approach is to compute the integral P (p R , pB ) = d p P (p | u) δ(pR − (p1 + p2 )) δ(pB − (p3 + p4 + p5 + p6 )). A cheaper approach is to compute the predictive distributions, given arbitrary data (F1 , F2 , F3 , F4 , F5 , F6 ), and find the condition for the two predictive distributions to match for all data.

The entropic distribution for a probability vector p is sometimes used in the ‘maximum entropy’ image reconstruction community. P 1 (23.35) exp[−αDKL (p||m)] δ( i pi − 1) , Z(α, m) P where m, the measure, is a positive vector, and D KL (p||m) = i pi log pi /mi . P (p | α, m) =

Further reading

See (MacKay and Peto, 1995) for fun with Dirichlets.

23.6 Further exercises Exercise 23.4.[2 ] N datapoints {xn } are drawn from a gamma distribution P (x | s, c) = Γ(x; s, c) with unknown parameters s and c. What are the maximum likelihood parameters s and c?

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24 Exact Marginalization

How can we avoid the exponentially large cost of complete enumeration of all hypotheses? Before we stoop to approximate methods, we explore two approaches to exact marginalization: first, marginalization over continuous variables (sometimes known as nuisance parameters) by doing integrals; and second, summation over discrete variables by message-passing. Exact marginalization over continuous parameters is a macho activity enjoyed by those who are fluent in definite integration. This chapter uses gamma distributions; as was explained in the previous chapter, gamma distributions are a lot like Gaussian distributions, except that whereas the Gaussian goes from −∞ to ∞, gamma distributions go from 0 to ∞.

24.1 Inferring the mean and variance of a Gaussian distribution We discuss again the one-dimensional Gaussian distribution, parameterized by a mean µ and a standard deviation σ:   (x − µ)2 1 exp − P (x | µ, σ) = √ ≡ Normal(x; µ, σ 2 ). 2σ 2 2πσ

(24.1)

When inferring these parameters, we must specify their prior distribution. The prior gives us the opportunity to include specific knowledge that we have about µ and σ (from independent experiments, or on theoretical grounds, for example). If we have no such knowledge, then we can construct an appropriate prior that embodies our supposed ignorance. In section 21.2, we assumed a uniform prior over the range of parameters plotted. If we wish to be able to perform exact marginalizations, it may be useful to consider conjugate priors; these are priors whose functional form combines naturally with the likelihood such that the inferences have a convenient form.

Conjugate priors for µ and σ The conjugate prior for a mean µ is a Gaussian: we introduce two ‘hyperparameters’, µ0 and σµ , which parameterize the prior on µ, and write P (µ | µ0 , σµ ) = Normal(µ; µ0 , σµ2 ). In the limit µ0 = 0, σµ → ∞, we obtain the noninformative prior for a location parameter, the flat prior. This is noninformative because it is invariant under the natural reparameterization µ0 = µ + c. The prior P (µ) = const. is also an improper prior, that is, it is not normalizable. The conjugate prior for a standard deviation σ is a gamma distribution, which has two parameters bβ and cβ . It is most convenient to define the prior 319

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320

24 — Exact Marginalization

density of the inverse variance (the precision parameter) β = 1/σ 2 : P (β) = Γ(β; bβ , cβ ) =

  β 1 β cβ −1 exp − , Γ(cβ ) bcββ bβ

0 ≤ β < ∞.

(24.2)

This is a simple peaked distribution with mean b β cβ and variance b2β cβ . In the limit bβ cβ = 1, cβ → 0, we obtain the noninformative prior for a scale parameter, the 1/σ prior. This is ‘noninformative’ because it is invariant under the reparameterization σ 0 = cσ. The 1/σ prior is less strange-looking if we examine the resulting density over ln σ, or ln β, which is flat. This is the prior that expresses ignorance about σ by saying ‘well, it could be 10, or it could be 1, or it could be 0.1, . . . ’ Scale variables such as σ are usually best represented in terms of their logarithm. Again, this noninformative 1/σ prior is improper. In the following examples, I will use the improper noninformative priors for µ and σ. Using improper priors is viewed as distasteful in some circles, so let me excuse myself by saying it’s for the sake of readability; if I included proper priors, the calculations could still be done but the key points would be obscured by the flood of extra parameters.

Maximum likelihood and marginalization: σN and σN−1 The task of inferring the mean and standard deviation of a Gaussian distribution from N samples is a familiar one, though maybe not everyone understands the difference between the σN and σN−1 buttons on their calculator. Let us recap the formulae, then derive them. Given data D = {xn }N n=1 , an ‘estimator’ of µ is x ¯≡

PN

n=1 xn /N,

and two estimators of σ are: s s PN PN 2 (x − x ¯ ) ¯ )2 n n=1 n=1 (xn − x σN ≡ and σN−1 ≡ . N N −1

(24.3)

(24.4)

There are two principal paradigms for statistics: sampling theory and Bayesian inference. In sampling theory (also known as ‘frequentist’ or orthodox statistics), one invents estimators of quantities of interest and then chooses between those estimators using some criterion measuring their sampling properties; there is no clear principle for deciding which criterion to use to measure the performance of an estimator; nor, for most criteria, is there any systematic procedure for the construction of optimal estimators. In Bayesian inference, in contrast, once we have made explicit all our assumptions about the model and the data, our inferences are mechanical. Whatever question we wish to pose, the rules of probability theory give a unique answer which consistently takes into account all the given information. Human-designed estimators and confidence intervals have no role in Bayesian inference; human input only enters into the important tasks of designing the hypothesis space (that is, the specification of the model and all its probability distributions), and figuring out how to do the computations that implement inference in that space. The answers to our questions are probability distributions over the quantities of interest. We often find that the estimators of sampling theory emerge automatically as modes or means of these posterior distributions when we choose a simple hypothesis space and turn the handle of Bayesian inference.

Reminder: when we change variables from σ to l(σ), a one-to-one function of σ, the probability density transforms from Pσ (σ) to ∂σ Pl (l) = Pσ (σ) . ∂l Here, the Jacobian is ∂σ ∂ ln σ = σ.

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24.1: Inferring the mean and variance of a Gaussian distribution

321

0.06 1 0.05

0.9

0.04

0.8

0.03

0.7 0.6

0.02

0.5 0.01

0.4

10

0.3

sigma

0.2

0.8 0.6 sigma

0.1 0

0.4

(a1)

0.2 0

0.5

1.5

1 mean

2

0.5

1

(a2)

0.09

1.5

2

mean

0.09 mu=1 mu=1.25 mu=1.5

0.08

0.08 P(sigma|D,mu=1)

0.07

0.07

0.06

0.06

0.05

0.05

0.04

0.04

0.03

0.03

P(sigma|D)

0.02

(c)

0.02

(d)

0.01 0 0.2

0.4

0.6

0.8

1

1.2 1.4 1.6 1.8 2

0.01 0 0.2

0.4

0.6

0.8

1

1.2 1.4 1.6 1.8 2

In sampling theory, the estimators above can be motivated as follows. x ¯ is an unbiased estimator of µ which, out of all the possible unbiased estimators of µ, has smallest variance (where this variance is computed by averaging over an ensemble of imaginary experiments in which the data samples are assumed to come from an unknown Gaussian distribution). The estimator (¯ x, σ N ) is the maximum likelihood estimator for (µ, σ). The estimator σ N is biased, however: the expectation of σN , given σ, averaging over many imagined experiments, is not σ. Exercise 24.1.[2, p.323] Give an intuitive explanation why the estimator σ N is biased. This bias motivates the invention, in sampling theory, of σ N−1 , which can be 2 shown to be an unbiased estimator. Or to be precise, it is σ N− 1 that is an unbiased estimator of σ 2 . We now look at some Bayesian inferences for this problem, assuming noninformative priors for µ and σ. The emphasis is thus not on the priors, but rather on (a) the likelihood function, and (b) the concept of marginalization. The joint posterior probability of µ and σ is proportional to the likelihood function illustrated by a contour plot in figure 24.1a. The log likelihood is: X √ ln P ({xn }N (xn − µ)2 /(2σ 2 ), (24.5) n=1 | µ, σ) = −N ln( 2πσ) − n

√ ¯)2 + S]/(2σ 2 ), (24.6) = −N ln( 2πσ) − [N (µ − x

P where S ≡ n (xn − x ¯)2 . Given the Gaussian model, the likelihood can be expressed in terms of the two functions of the data x ¯ and S, so these two quantities are known as ‘sufficient statistics’. The posterior probability of µ and σ is, using the improper priors: P (µ, σ | {xn }N n=1 ) = =

P ({xn }N n=1 | µ, σ)P (µ, σ) P ({xn }N ) n=1  N (µ−¯ x)2 +S 1 1 1 exp − 2 2 N/2 σµ σ 2σ (2πσ ) P ({xn }N n=1 )

(24.7)

.

(24.8)

Figure 24.1. The likelihood function for the parameters of a Gaussian distribution, repeated from figure 21.5. (a1, a2) Surface plot and contour plot of the log likelihood as a function of µ and σ. The data set of N = 5 P points had mean x ¯ = 1.0 and S = (x − x ¯)2 = 1.0. Notice that the maximum is skew in σ. The two estimators of standard deviation have values σN = 0.45 and σN−1 = 0.50. (c) The posterior probability of σ for various fixed values of µ (shown as a density over ln σ). (d) The posterior probability of σ, P (σ | D), assuming a flat prior on µ, obtained by projecting the probability mass in (a) onto the σ axis. The maximum of P (σ | D) is at σN−1 . By contrast, the maximum of P (σ | D, µ = x ¯) is at σN . (Both probabilities are shows as densities over ln σ.)

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322

24 — Exact Marginalization

This function describes the answer to the question, ‘given the data, and the noninformative priors, what might µ and σ be?’ It may be of interest to find the parameter values that maximize the posterior probability, though it should be emphasized that posterior probability maxima have no fundamental status in Bayesian inference, since their location depends on the choice of basis. Here we choose the basis (µ, ln σ), in which our prior is flat, so that the posterior probability maximum coincides with the maximum of the likelihood. As we saw in exercisen 22.4 (p.302), the o maximum likelihood solution for µ and ln σ p is {µ, σ}ML = x ¯, σN = S/N . There is more to the posterior distribution than just its mode. As can be seen in figure 24.1a, the likelihood has a skew peak. As we increase σ, the width of the conditional distribution of µ increases (figure 22.1b). And if we fix µ to a sequence of values moving away from the sample mean x ¯, we obtain a sequence of conditional distributions over σ whose maxima move to increasing values of σ (figure 24.1c). The posterior probability of µ given σ is P (µ | {xn }N n=1 , σ) =

P ({xn }N n=1 | µ, σ)P (µ) P ({xn }N n=1 | σ)

∝ exp(−N (µ − x ¯)2 /(2σ 2 )) = Normal(µ; x¯, σ 2 /N ).

(24.9) (24.10) (24.11)

√ We note the familiar σ/ N scaling of the error bars on µ. Let us now ask the question ‘given the data, and the noninformative priors, what might σ be?’ This question differs from the first one we asked in that we are now not interested in µ. This parameter must therefore be marginalized over. The posterior probability of σ is: P (σ | {xn }N n=1 ) =

P ({xn }N n=1 | σ)P (σ) . P ({xn }N n=1 )

(24.12)

The data-dependent term P ({xn }N n=1 | σ) appeared earlier as the normalizing constant in equation (24.9); one name for this quantity is the ‘evidence’, or marginal likelihood, for σ. We obtain the evidence for σ by integrating out µ; a noninformative prior P (µ) = constant is assumed; we call this constant 1/σµ , so that we can think of the prior prior of width σ µ . The R as a top-hat Gaussian integral, P ({xn }N P ({xn }N n=1 | σ) = n=1 | µ, σ)P (µ) dµ, yields: √ √ √ 2πσ/ N S N . (24.13) ln P ({xn }n=1 | σ) = −N ln( 2πσ) − 2 + ln 2σ σµ The first two terms are the best-fit log likelihood (i.e., the log likelihood with µ=x ¯). The last term is the log of the Occam factor which penalizes smaller values of σ. (We will discuss Occam factors more in Chapter 28.) When we differentiate the log evidence with respect to ln σ, to find the most probable √ σ, the additional volume factor (σ/ N ) shifts the maximum from σN to p σN−1 = S/(N − 1). (24.14)

Intuitively, the denominator (N −1) number of noise measurements P counts the contained in the quantity S = n (xn − x ¯)2 . The sum contains N residuals squared, but there are only (N −1) effective noise measurements because the determination of one parameter µ from the data causes one dimension of noise to be gobbled up in unavoidable overfitting. In the terminology of classical

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24.2: Exercises

323

statistics, the Bayesian’s best guess for σ sets χ 2 (the measure of deviance P 2 ˆ )2 /ˆ σ 2 ) equal to the number of degrees of freedom, defined by χ ≡ n (xn − µ N − 1. Figure 24.1d shows the posterior probability of σ, which is proportional to the marginal likelihood. This may be contrasted with the posterior probability of σ with µ fixed to its most probable value, x ¯ = 1, which is shown in figure 24.1c and d. The final inference we might wish to make is ‘given the data, what is µ?’ . Exercise 24.2.[3 ] Marginalize over σ and obtain the posterior marginal distribution of µ, which is a Student-t distribution: P (µ | D) ∝ 1/ N (µ − x ¯ )2 + S

Further reading

N/2

.

(24.15)

A bible of exact marginalization is Bretthorst’s (1988) book on Bayesian spectrum analysis and parameter estimation.

24.2 Exercises . Exercise 24.3.[3 ] [This exercise requires macho integration capabilities.] Give a Bayesian solution to exercise 22.15 (p.309), where seven scientists of varying capabilities have measured µ with personal noise levels σ n , and we are interested in inferring µ. Let the prior on each σ n be a broad prior, for example a gamma distribution with parameters (s, c) = (10, 0.1). Find the posterior distribution of µ. Plot it, and explore its properties for a variety of data sets such as the one given, and the data set {xn } = {13.01, 7.39}.

[Hint: first find the posterior distribution of σ n given µ and xn , P (σn | xn , µ). Note that the normalizing constant for this inference is P (xn | µ). Marginalize over σn to find this normalizing constant, then use Bayes’ theorem a second time to find P (µ | {x n }).]

24.3 Solutions Solution to exercise 24.1 (p.321). 1. The data points are distributed with mean squared deviation σ 2 about the true mean. 2. The sample mean is unlikely to exactly equal the true mean. 3. The sample mean is the value of µ that minimizes the sum squared deviation of the data points from µ. Any other value of µ (in particular, the true value of µ) will have a larger value of the sum-squared deviation that µ = x ¯. So the expected mean squared deviation from the sample mean is necessarily smaller than the mean squared deviation σ 2 about the true mean.

A

-30

B C D-G

-20

-10

0

10

20

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25 Exact Marginalization in Trellises In this chapter we will discuss a few exact methods that are used in probabilistic modelling. As an example we will discuss the task of decoding a linear error-correcting code. We will see that inferences can be conducted most efficiently by message-passing algorithms, which take advantage of the graphical structure of the problem to avoid unnecessary duplication of computations (see Chapter 16).

25.1 Decoding problems A codeword t is selected from a linear (N, K) code C, and it is transmitted over a noisy channel; the received signal is y. In this chapter we will assume that the channel is a memoryless channel such as a Gaussian channel. Given an assumed channel model P (y | t), there are two decoding problems. The codeword decoding problem is the task of inferring which codeword t was transmitted given the received signal. The bitwise decoding problem is the task of inferring for each transmitted bit tn how likely it is that that bit was a one rather than a zero. As a concrete example, take the (7, 4) Hamming code. In Chapter 1, we discussed the codeword decoding problem for that code, assuming a binary symmetric channel. We didn’t discuss the bitwise decoding problem and we didn’t discuss how to handle more general channel models such as a Gaussian channel.

Solving the codeword decoding problem By Bayes’ theorem, the posterior probability of the codeword t is P (t | y) =

P (y | t)P (t) . P (y)

(25.1)

Likelihood function. The first factor in the numerator, P (y | t), is the likelihood of the codeword, which, for any memoryless channel, is a separable function, N Y P (y | t) = P (yn | tn ). (25.2) n=1

For example, if the channel is a Gaussian channel with transmissions ±x and additive noise of standard deviation σ, then the probability density 324

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25.1: Decoding problems

325

of the received signal yn in the two cases tn = 0, 1 is   1 (yn − x)2 P (yn | tn = 1) = √ exp − 2σ 2 2πσ 2   (yn + x)2 1 exp − . P (yn | tn = 0) = √ 2σ 2 2πσ 2

(25.3) (25.4)

From the point of view of decoding, all that matters is the likelihood ratio, which for the case of the Gaussian channel is   P (yn | tn = 1) 2xyn = exp . (25.5) P (yn | tn = 0) σ2 Exercise 25.1.[2 ] Show that from the point of view of decoding, a Gaussian channel is equivalent to a time-varying binary symmetric channel with a known noise level fn which depends on n. Prior. The second factor in the numerator is the prior probability of the codeword, P (t), which is usually assumed to be uniform over all valid codewords. The denominator in (25.1) is the normalizing constant P (y) =

X t

P (y | t)P (t).

(25.6)

The complete solution to the codeword decoding problem is a list of all codewords and their probabilities as given by equation (25.1). Since the number of codewords in a linear code, 2K , is often very large, and since we are not interested in knowing the detailed probabilities of all the codewords, we often restrict attention to a simplified version of the codeword decoding problem. The MAP codeword decoding problem is the task of identifying the most probable codeword t given the received signal. If the prior probability over codewords is uniform then this task is identical to the problem of maximum likelihood decoding, that is, identifying the codeword that maximizes P (y | t). Example: In Chapter 1, for the (7, 4) Hamming code and a binary symmetric channel we discussed a method for deducing the most probable codeword from the syndrome of the received signal, thus solving the MAP codeword decoding problem for that case. We would like a more general solution. The MAP codeword decoding problem can be solved in exponential time (of order 2K ) by searching through all codewords for the one that maximizes P (y | t)P (t). But we are interested in methods that are more efficient than this. In section 25.3, we will discuss an exact method known as the min–sum algorithm which may be able to solve the codeword decoding problem more efficiently; how much more efficiently depends on the properties of the code. It is worth emphasizing that MAP codeword decoding for a general linear code is known to be NP-complete (which means in layman’s terms that MAP codeword decoding has a complexity that scales exponentially with the blocklength, unless there is a revolution in computer science). So restricting attention to the MAP decoding problem hasn’t necessarily made the task much less challenging; it simply makes the answer briefer to report.

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326

25 — Exact Marginalization in Trellises

Solving the bitwise decoding problem Formally, the exact solution of the bitwise decoding problem is obtained from equation (25.1) by marginalizing over the other bits. X P (tn | y) = P (t | y). (25.7) {tn0 : n0 6=n}

We can also write this marginal with the aid of a truth function [S] that is one if the proposition S is true and zero otherwise. X P (tn = 1 | y) = (25.8) P (t | y) [tn = 1] 



t

P (tn = 0 | y) =

X t

P (t | y) [tn = 0]. 

(25.9)

Computing these marginal probabilities by an explicit sum over all codewords t takes exponential time. But, for certain codes, the bitwise decoding problem can be solved much more efficiently using the forward–backward algorithm. We will describe this algorithm, which is an example of the sum–product algorithm, in a moment. Both the min–sum algorithm and the sum–product algorithm have widespread importance, and have been invented many times in many fields.

Repetition code R3 (b)

25.2 Codes and trellises In Chapters 1 and 11, we represented linear (N, K) codes in terms of their generator matrices and their parity-check matrices. In the case of a systematic block code, the first K transmitted bits in each block of size N are the source bits, and the remaining M = N −K bits are the parity-check bits. This means that the generator matrix of the code can be written   IK , (25.10) GT = P and the parity-check matrix can be written   H = P IM ,

(a)

Simple parity code P3

(c)

(25.11)

where P is an M × K matrix. In this section we will study another representation of a linear code called a trellis. The codes that these trellises represent will not in general be systematic codes, but they can be mapped onto systematic codes if desired by a reordering of the bits in a block.

Definition of a trellis Our definition will be quite narrow. For a more comprehensive view of trellises, the reader should consult Kschischang and Sorokine (1995). A trellis is a graph consisting of nodes (also known as states or vertices) and edges. The nodes are grouped into vertical slices called times, and the times are ordered such that each edge connects a node in one time to a node in a neighbouring time. Every edge is labelled with a symbol. The leftmost and rightmost states contain only one node. Apart from these two extreme nodes, all nodes in the trellis have at least one edge connecting leftwards and at least one connecting rightwards.

(7, 4) Hamming code

Figure 25.1. Examples of trellises. Each edge in a trellis is labelled by a zero (shown by a square) or a one (shown by a cross).

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25.3: Solving the decoding problems on a trellis A trellis with N + 1 times defines a code of blocklength N as follows: a codeword is obtained by taking a path that crosses the trellis from left to right and reading out the symbols on the edges that are traversed. Each valid path through the trellis defines a codeword. We will number the leftmost time ‘time 0’ and the rightmost ‘time N ’. We will number the leftmost state ‘state 0’ and the rightmost ‘state I’, where I is the total number of states (vertices) in the trellis. The nth bit of the codeword is emitted as we move from time n−1 to time n. The width of the trellis at a given time is the number of nodes in that time. The maximal width of a trellis is what it sounds like. A trellis is called a linear trellis if the code it defines is a linear code. We will solely be concerned with linear trellises from now on, as nonlinear trellises are much more complex beasts. For brevity, we will only discuss binary trellises, that is, trellises whose edges are labelled with zeroes and ones. It is not hard to generalize the methods that follow to q-ary trellises. Figures 25.1(a–c) show the trellises corresponding to the repetition code R3 which has (N, K) = (3, 1); the parity code P 3 with (N, K) = (3, 2); and the (7, 4) Hamming code. . Exercise 25.2.[2 ] Confirm that the sixteen codewords listed in table 1.14 are generated by the trellis shown in figure 25.1c.

Observations about linear trellises For any linear code the minimal trellis is the one that has the smallest number of nodes. In a minimal trellis, each node has at most two edges entering it and at most two edges leaving it. All nodes in a time have the same left degree as each other and they have the same right degree as each other. The width is always a power of two. A minimal trellis for a linear (N, K) code cannot have a width greater than 2K since every node has at least one valid codeword through it, and there are only 2K codewords. Furthermore, if we define M = N − K, the minimal trellis’s width is everywhere less than 2 M . This will be proved in section 25.4. Notice that for the linear trellises in figure 25.1, all of which are minimal trellises, K is the number of times a binary branch point is encountered as the trellis is traversed from left to right or from right to left. We will discuss the construction of trellises more in section 25.4. But we now know enough to discuss the decoding problem.

25.3 Solving the decoding problems on a trellis We can view the trellis of a linear code as giving a causal description of the probabilistic process that gives rise to a codeword, with time flowing from left to right. Each time a divergence is encountered, a random source (the source of information bits for communication) determines which way we go. At the receiving end, we receive a noisy version of the sequence of edgelabels, and wish to infer which path was taken, or to be precise, (a) we want to identify the most probable path in order to solve the codeword decoding problem; and (b) we want to find the probability that the transmitted symbol at time n was a zero or a one, to solve the bitwise decoding problem. Example 25.3. Consider the case of a single transmission from the Hamming (7, 4) trellis shown in figure 25.1c.

327

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328

25 — Exact Marginalization in Trellises t

Likelihood

Posterior probability

0000000 0001011 0010111 0011100 0100110 0101101 0110001 0111010 1000101 1001110 1010010 1011001 1100011 1101000 1110100 1111111

0.0275562 0.0001458 0.0013122 0.0030618 0.0002268 0.0000972 0.0708588 0.0020412 0.0001458 0.0000042 0.0030618 0.0013122 0.0000972 0.0002268 0.0020412 0.0000108

0.25 0.0013 0.012 0.027 0.0020 0.0009 0.63 0.018 0.0013 0.0000 0.027 0.012 0.0009 0.0020 0.018 0.0001

Figure 25.2. Posterior probabilities over the sixteen codewords when the received vector y has normalized likelihoods (0.1, 0.4, 0.9, 0.1, 0.1, 0.1, 0.3).

Let the normalized likelihoods be: (0.1, 0.4, 0.9, 0.1, 0.1, 0.1, 0.3). That is, the ratios of the likelihoods are P (y1 | x1 = 1) 0.1 P (y2 | x2 = 1) 0.4 = , = , etc. P (y1 | x1 = 0) 0.9 P (y2 | x2 = 0) 0.6

(25.12)

How should this received signal be decoded? 1. If we threshold the likelihoods at 0.5 to turn the signal into a binary received vector, we have r = (0, 0, 1, 0, 0, 0, 0), which decodes, using the decoder for the binary symmetric channel (Chapter 1), into ˆt = (0, 0, 0, 0, 0, 0, 0). This is not the optimal decoding procedure. Optimal inferences are always obtained by using Bayes’ theorem. 2. We can find the posterior probability over codewords by explicit enumeration of all sixteen codewords. This posterior distribution is shown in figure 25.2. Of course, we aren’t really interested in such brute-force solutions, and the aim of this chapter is to understand algorithms for getting the same information out in less than 2 K computer time. Examining the posterior probabilities, we notice that the most probable codeword is actually the string t = 0110001. This is more than twice as probable as the answer found by thresholding, 0000000. Using the posterior probabilities shown in figure 25.2, we can also compute the posterior marginal distributions of each of the bits. The result is shown in figure 25.3. Notice that bits 1, 4, 5 and 6 are all quite confidently inferred to be zero. The strengths of the posterior probabilities for bits 2, 3, and 7 are not so great. 2 In the above example, the MAP codeword is in agreement with the bitwise decoding that is obtained by selecting the most probable state for each bit using the posterior marginal distributions. But this is not always the case, as the following exercise shows.

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25.3: Solving the decoding problems on a trellis

n 1 2 3 4 5 6 7

Likelihood P (yn | tn = 1) P (yn | tn = 0) 0.1 0.4 0.9 0.1 0.1 0.1 0.3

0.9 0.6 0.1 0.9 0.9 0.9 0.7

329

Posterior marginals P (tn = 1 | y) P (tn = 0 | y) 0.061 0.674 0.746 0.061 0.061 0.061 0.659

0.939 0.326 0.254 0.939 0.939 0.939 0.341

Exercise 25.4.[2, p.333] Find the most probable codeword in the case where the normalized likelihood is (0.2, 0.2, 0.9, 0.2, 0.2, 0.2, 0.2). Also find or estimate the marginal posterior probability for each of the seven bits, and give the bit-by-bit decoding. [Hint: concentrate on the few codewords that have the largest probability.] We now discuss how to use message passing on a code’s trellis to solve the decoding problems.

The min–sum algorithm The MAP codeword decoding problem can be solved using the min–sum algorithm that was introduced in section 16.3. Each codeword of the code corresponds to a path across the trellis. Just as the cost of a journey is the sum of the costs of its constituent steps, the log likelihood of a codeword is the sum of the bitwise log likelihoods. By convention, we flip the sign of the log likelihood (which we would like to maximize) and talk in terms of a cost, which we would like to minimize. We associate with each edge a cost −log P (y n | tn ), where tn is the transmitted bit associated with that edge, and y n is the received symbol. The min–sum algorithm presented in section 16.3 can then identify the most probable codeword in a number of computer operations equal to the number of edges in the trellis. This algorithm is also known as the Viterbi algorithm (Viterbi, 1967).

The sum–product algorithm To solve the bitwise decoding problem, we can make a small modification to the min–sum algorithm, so that the messages passed through the trellis define ‘the probability of the data up to the current point’ instead of ‘the cost of the best route to this point’. We replace the costs on the edges, −log P (y n | tn ), by the likelihoods themselves, P (yn | tn ). We replace the min and sum operations of the min–sum algorithm by a sum and product respectively. Let i run over nodes/states, i = 0 be the label for the start state, P(i) denote the set of states that are parents of state i, and w ij be the likelihood associated with the edge from node j to node i. We define the forward-pass messages αi by α0 = 1

Figure 25.3. Marginal posterior probabilities for the 7 bits under the posterior distribution of figure 25.2.

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330

25 — Exact Marginalization in Trellises X

αi =

wij αj .

(25.13)

j∈P(i)

These messages can be computed sequentially from left to right. . Exercise 25.5.[2 ] Show that for a node i whose time-coordinate is n, α i is proportional to the joint probability that the codeword’s path passed through node i and that the first n received symbols were y 1 , . . . , yn . The message αI computed at the end node of the trellis is proportional to the marginal probability of the data. . Exercise 25.6.[2 ] What is the constant of proportionality? [Answer: 2 K ] We define a second set of backward-pass messages β i in a similar manner. Let node I be the end node. βI

= 1

βj

=

X

wij βi .

(25.14)

i:j∈P(i)

These messages can be computed sequentially in a backward pass from right to left. . Exercise 25.7.[2 ] Show that for a node i whose time-coordinate is n, β i is proportional to the conditional probability, given that the codeword’s path passed through node i, that the subsequent received symbols were yn+1 . . . yN . Finally, to find the probability that the nth bit was a 1 or 0, we do two summations of products of the forward and backward messages. Let i run over nodes at time n and j run over nodes at time n − 1, and let t ij be the value of tn associated with the trellis edge from node j to node i. For each value of t = 0/1, we compute X rn(t) = αj wij βi . (25.15) i,j: j∈P(i), tij =t

Then the posterior probability that t n was t = 0/1 is P (tn = t | y) = (0)

1 (t) r , Z n

(25.16)

(1)

where the normalizing constant Z = r n + rn should be identical to the final forward message αI that was computed earlier.

n

[2 ]

Exercise 25.8. Confirm that the above sum–product algorithm does compute P (tn = t | y). Other names for the sum–product algorithm presented here are ‘the forward– backward algorithm’, ‘the BCJR algorithm’, and ‘belief propagation’. . Exercise 25.9.[2, p.333] A codeword of the simple parity code P 3 is transmitted, and the received signal y has associated likelihoods shown in table 25.4. Use the min–sum algorithm and the sum–product algorithm in the trellis (figure 25.1) to solve the MAP codeword decoding problem and the bitwise decoding problem. Confirm your answers by enumeration of all codewords (000, 011, 110, 101). [Hint: use logs to base 2 and do the min–sum computations by hand. When working the sum–product algorithm by hand, you may find it helpful to use three colours of pen, one for the αs, one for the ws, and one for the βs.]

1 2 3

P (yn | tn ) tn = 0 t n = 1 1/4 1/2

1/2 1/4

1/8

1/2

Table 25.4. Bitwise likelihoods for a codeword of P3 .

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25.4: More on trellises

331

25.4 More on trellises We now discuss various ways of making the trellis of a code. You may safely jump over this section. The span of a codeword is the set of bits contained between the first bit in the codeword that is non-zero, and the last bit that is non-zero, inclusive. We can indicate the span of a codeword by a binary vector as shown in table 25.5.

Codeword Span

0000000 0000000

0001011 0001111

0100110 0111110

1100011 1111111

0101101 0111111

A generator matrix is in trellis-oriented form if the spans of the rows of the generator matrix all start in different columns and the spans all end in different columns.

How to make a trellis from a generator matrix First, put the generator matrix into trellis-oriented form by row-manipulations similar to Gaussian elimination. For example, our (7, 4) Hamming code can be generated by   1 0 0 0 1 0 1  0 1 0 0 1 1 0   (25.17) G=  0 0 1 0 1 1 1  0 0 0 1 0 1 1 but this matrix is not in trellis-oriented form – for example, rows 1, 3 and 4 all have spans that end in the same column. By subtracting lower rows from upper rows, we can obtain an equivalent generator matrix (that is, one that generates the same set of codewords) as follows: 

1  0 G=  0 0

1 1 0 0

0 0 1 0

1 0 1 1

0 1 1 0

0 1 0 1

 0 0  . 0  1

(25.18)

Now, each row of the generator matrix can be thought of as defining an (N, 1) subcode of the (N, K) code, that is, in this case, a code with two codewords of length N = 7. For the first row, the code consists of the two codewords 1101000 and 0000000. The subcode defined by the second row consists of 0100110 and 0000000. It is easy to construct the minimal trellises of these subcodes; they are shown in the left column of figure 25.6. We build the trellis incrementally as shown in figure 25.6. We start with the trellis corresponding to the subcode given by the first row of the generator matrix. Then we add in one subcode at a time. The vertices within the span of the new subcode are all duplicated. The edge symbols in the original trellis are left unchanged and the edge symbols in the second part of the trellis are flipped wherever the new subcode has a 1 and otherwise left alone. Another (7, 4) Hamming code can be generated by 

1  0 G=  0 0

1 1 0 0

1 1 1 0

0 1 0 1

0 1 1 1

0 0 1 1

 0 0  . 0  1

(25.19)

Table 25.5. Some codewords and their spans.

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332

25 — Exact Marginalization in Trellises Figure 25.6. Trellises for four subcodes of the (7, 4) Hamming code (left column), and the sequence of trellises that are made when constructing the trellis for the (7, 4) Hamming code (right column). Each edge in a trellis is labelled by a zero (shown by a square) or a one (shown by a cross).

+ =

+

=

+ =

The (7, 4) Hamming code generated by this matrix differs by a permutation of its bits from the code generated by the systematic matrix used in Chapter 1 and above. The parity-check matrix corresponding to this permutation is: 

 1 0 1 0 1 0 1 H =  0 1 1 0 0 1 1 . 0 0 0 1 1 1 1

(25.20)

(a)

The trellis obtained from the permuted matrix G given in equation (25.19) is shown in figure 25.7a. Notice that the number of nodes in this trellis is smaller than the number of nodes in the previous trellis for the Hamming (7, 4) code in figure 25.1c. We thus observe that rearranging the order of the codeword bits can sometimes lead to smaller, simpler trellises. (b)

Trellises from parity-check matrices Another way of viewing the trellis is in terms of the syndrome. The syndrome of a vector r is defined to be Hr, where H is the parity-check matrix. A vector is only a codeword if its syndrome is zero. As we generate a codeword we can describe the current state by the partial syndrome, that is, the product of H with the codeword bits thus far generated. Each state in the trellis is a partial syndrome at one time coordinate. The starting and ending states are both constrained to be the zero syndrome. Each node in a state represents a different possible value for the partial syndrome. Since H is an M × N matrix, where M = N − K, the syndrome is at most an M -bit vector. So we need at most 2M nodes in each state. We can construct the trellis of a code from its parity-check matrix by walking from each end, generating two trees of possible syndrome sequences. The intersection of these two trees defines the trellis of the code. In the pictures we obtain from this construction, we can let the vertical coordinate represent the syndrome. Then any horizontal edge is necessarily associated with a zero bit (since only a non-zero bit changes the syndrome)

Figure 25.7. Trellises for the permuted (7, 4) Hamming code generated from (a) the generator matrix by the method of figure 25.6; (b) the parity-check matrix by the method on page 332. Each edge in a trellis is labelled by a zero (shown by a square) or a one (shown by a cross).

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25.5: Solutions

333

and any non-horizontal edge is associated with a one bit. (Thus in this representation we no longer need to label the edges in the trellis.) Figure 25.7b shows the trellis corresponding to the parity-check matrix of equation (25.20).

25.5 Solutions

t 0000000 0001011 0010111 0011100 0100110 0101101 0110001 0111010 1000101 1001110 1010010 1011001 1100011 1101000 1110100 1111111

Likelihood

Posterior probability

0.026 0.00041 0.0037 0.015 0.00041 0.00010 0.015 0.0037 0.00041 0.00010 0.015 0.0037 0.00010 0.00041 0.0037 0.000058

0.3006 0.0047 0.0423 0.1691 0.0047 0.0012 0.1691 0.0423 0.0047 0.0012 0.1691 0.0423 0.0012 0.0047 0.0423 0.0007

Solution to exercise 25.4 (p.329). The posterior probability over codewords is shown in table 25.8. The most probable codeword is 0000000. The marginal posterior probabilities of all seven bits are: n 1 2 3 4 5 6 7

Likelihood P (yn | tn = 1) P (yn | tn = 0) 0.2 0.2 0.9 0.2 0.2 0.2 0.2

0.8 0.8 0.1 0.8 0.8 0.8 0.8

Posterior marginals P (tn = 1 | y) P (tn = 0 | y) 0.266 0.266 0.677 0.266 0.266 0.266 0.266

0.734 0.734 0.323 0.734 0.734 0.734 0.734

So the bitwise decoding is 0010000, which is not actually a codeword. Solution to exercise 25.9 (p.330). The MAP codeword is 101, and its likelihood is 1/8. The normalizing constant of the sum–product algorithm is Z = αI = 3/16. The intermediate αi are (from left to right) 1/2, 1/4, 5/16, 4/16; the intermediate βi are (from right to left), 1/2, 1/8, 9/32, 3/16. The bitwise decoding is: P (t1 = 1 | y) = 3/4; P (t1 = 1 | y) = 1/4; P (t1 = 1 | y) = 5/6. The codewords’ probabilities are 1/12, 2/12, 1/12, 8/12 for 000, 011, 110, 101.

Table 25.8. The posterior probability over codewords for exercise 25.4.

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26 Exact Marginalization in Graphs We now take a more general view of the tasks of inference and marginalization. Before reading this chapter, you should read about message passing in Chapter 16.

26.1 The general problem Assume that a function P ∗ of a set of N variables x ≡ {xn }N n=1 is defined as a product of M factors as follows: P ∗ (x) =

M Y

fm (xm ).

(26.1)

m=1

Each of the factors fm (xm ) is a function of a subset xm of the variables that make up x. If P ∗ is a positive function then we may be interested in a second normalized function, P (x) ≡

1 ∗ P (x) Z

=

1 Z

M Y

fm (xm ),

(26.2)

m=1

where the normalizing constant Z is defined by Z=

M XY

fm (xm ).

(26.3)

x m=1

As an example of the notation we’ve just introduced, here’s a function of three binary variables x1 , x2 , x3 defined by the five factors:  0.1 x1 = 0 f1 (x1 ) =  0.9 x1 = 1 0.1 x2 = 0 f2 (x2 ) =  0.9 x2 = 1 0.9 x3 = 0 f3 (x3 ) =  0.1 x3 = 1 (26.4) 1 (x1 , x2 ) = (0, 0) or (1, 1) f4 (x1 , x2 ) =  0 (x1 , x2 ) = (1, 0) or (0, 1) 1 (x2 , x3 ) = (0, 0) or (1, 1) f5 (x2 , x3 ) = 0 (x2 , x3 ) = (1, 0) or (0, 1) P ∗ (x) = f1 (x1 )f2 (x2 )f3 (x3 )f4 (x1 , x2 )f5 (x2 , x3 ) P (x) = Z1 f1 (x1 )f2 (x2 )f3 (x3 )f4 (x1 , x2 )f5 (x2 , x3 ).

The five subsets of {x1 , x2 , x3 } denoted by xm in the general function (26.1) are here x1 = {x1 }, x2 = {x2 }, x3 = {x3 }, x4 = {x1 , x2 }, and x5 = {x2 , x3 }. 334

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26.1: The general problem

335

The function P (x), by the way, may be recognized as the posterior probability distribution of the three transmitted bits in a repetition code (section 1.2) when the received signal is r = (1, 1, 0) and the channel is a binary symmetric channel with flip probability 0.1. The factors f 4 and f5 respectively enforce the constraints that x1 and x2 must be identical and that x2 and x3 must be identical. The factors f1 , f2 , f3 are the likelihood functions contributed by each component of r. A function of the factored form (26.1) can be depicted by a factor graph, in which the variables are depicted by circular nodes and the factors are depicted by square nodes. An edge is put between variable node n and factor node m if the function fm (xm ) has any dependence on variable xn . The factor graph for the example function (26.4) is shown in figure 26.1.

The normalization problem The first task to be solved is to compute the normalizing constant Z.

The marginalization problems The second task to be solved is to compute the marginal function of any variable xn , defined by X Zn (xn ) = P ∗ (x). (26.5) {xn0 }, n0 6=n

For example, if f is a function of three variables then the marginal for n = 1 is defined by X f (x1 , x2 , x3 ). (26.6) Z1 (x1 ) = x2 ,x3

This type of summation, over ‘all the x n0 except for xn’ is so important that it can be useful to have a special notation for it – the ‘not-sum’ or ‘summary’. The third task to be solved is to compute the normalized marginal of any variable xn , defined by X Pn (xn ) ≡ P (x). (26.7) {xn0 }, n0 6=n

[We include the suffix ‘n’ in Pn (xn ), departing from our normal practice in the rest of the book, where we would omit it.] . Exercise 26.1.[1 ] Show that the normalized marginal is related to the marginal Zn (xn ) by Zn (xn ) . (26.8) Pn (xn ) = Z We might also be interested in marginals over a subset of the variables, such as X Z12 (x1 , x2 ) ≡ P ∗ (x1 , x2 , x3 ). (26.9) x3

All these tasks are intractable in general. Even if every factor is a function of only three variables, the cost of computing exact solutions for Z and for the marginals is believed in general to grow exponentially with the number of variables N . For certain functions P ∗ , however, the marginals can be computed efficiently by exploiting the factorization of P ∗ . The idea of how this efficiency

g @ @

x1

f1

f2

f3

f4

g @ @

x2

f5

Figure 26.1. The factor graph associated with the function P ∗ (x) (26.4).

g

x3

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336

26 — Exact Marginalization in Graphs

arises is well illustrated by the message-passing examples of Chapter 16. The sum–product algorithm that we now review is a generalization of messagepassing rule-set B (p.242). As was the case there, the sum–product algorithm is only valid if the graph is tree-like.

26.2 The sum–product algorithm Notation We identify the set of variables that the mth factor depends on, x m , by the set of their indices N (m). For our example function (26.4), the sets are N (1) = {1} (since f1 is a function of x1 alone),