Source coding

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Last modified: 2026-11-01 19:38:09 (PST)

1 Codes

Definition 1 (Binary code) Let \(X\) be a discrete random variable. A binary code for \(X\) assigns each value \(x\) a codeword, a finite string of binary digits, of length \(\ell(x)\).

Definition 2 (Uniquely decodable code) A binary code is uniquely decodable if every finite sequence of codewords, written one after another, can be split back into codewords in only one way.

Definition 3 (Expected length) The expected length of a binary code for \(X\), where \(X\) has probability mass function \(p(x) \stackrel{\text{def}}{=}\mathbb{P}(X=x)\), is

\[L \stackrel{\text{def}}{=}{\mathbb{E}\mathopen{}\left[\ell(X)\right]\mathclose{}} = \sum_{x} p(x) \, \ell(x).\]

Remark 1 (Short codewords for common values). A code with a small expected length gives frequent values short codewords and rare values long ones, since each value’s length counts in \(L\) in proportion to its probability.

2 The source coding theorem

Theorem 1 (Source coding theorem) Let \(X\) be a discrete random variable that takes finitely many values, at least two of them with positive probability, and let \(H_2(X)\) be its entropy with base-2 logarithms.

Every uniquely decodable binary code for \(X\) has expected length \(L \ge H_2(X)\).
Some uniquely decodable binary code for \(X\) has expected length \(L < H_2(X) + 1\).
Coding blocks of \(n\) independent copies of \(X\) at a time, some uniquely decodable binary code spends, on average, fewer than \(H_2(X) + 1/n\) binary digits per copy.

Remark 2 (Where the theorem comes from). Shannon (1948) proved that entropy is the limit on how far the output of a source can be compressed. Cover and Thomas (2006, sec. 5.3 to 5.5) prove each part of Theorem 1. For part 2, they give each value \(x\) with \(p(x) > 0\) the length \(\ell(x) = \lceil -\log_2 p(x) \rceil\), its surprisal in bits rounded up; the Kraft inequality (Cover and Thomas 2006, sec. 5.2) guarantees that a code with these lengths exists.

3 Coding with the wrong distribution

Theorem 2 (The cost of coding with the wrong distribution) Let \(X\) take finitely many values, with probability mass function \(p\), and let \(q\) be a probability mass function, possibly different from \(p\), with \(q(x) > 0\) wherever \(p(x) > 0\). Suppose a code gives each value \(x\) with \(p(x) > 0\) the length \(\ell(x) = \lceil -\log_2 q(x) \rceil\). Then its expected length \(L\) under \(p\) satisfies

\[H_2(p, q) \le L < H_2(p, q) + 1,\]

where \(H_2(p, q)\) is the cross-entropy of \(q\) relative to \(p\), with base-2 logarithms.

Proof

Proof. Every sum here runs over the values \(x\) with \(p(x) > 0\). Rounding up adds less than one binary digit, so for each such \(x\)

\[-\log_2 q(x) \le \ell(x) < -\log_2 q(x) + 1.\]

Multiply each part by \(p(x) > 0\) and sum. Since \(\sum_{x} p(x) = 1\),

\[ -\sum_{x} p(x) \log_2 q(x) \le \sum_{x} p(x) \, \ell(x) < -\sum_{x} p(x) \log_2 q(x) + 1. \]

The middle sum is \(L\), and the outer sums are \(H_2(p, q)\).

Remark 3 (Cross-entropy is the cost of the wrong code). Write \(H_2(p)\) for the entropy of a random variable with mass function \(p\), and \(D_2(p \,\|\, q)\) for the Kullback–Leibler divergence of \(q\) from \(p\), both with base-2 logarithms. Cross-entropy splits into those two parts:

\[H_2(p, q) = H_2(p) + D_2(p \,\|\, q).\]

By Theorem 1, no uniquely decodable code averages fewer than \(H_2(p)\) bits. By Theorem 2, a code built for \(q\) costs about \(D_2(p \,\|\, q)\) bits per value more than that. The divergence is never negative, and it is \(0\) only when \(q = p\) (Cover and Thomas 2006, sec. 2.6), so the excess vanishes only when the code was built for the true distribution. Cover and Thomas (2006, sec. 5.4) state Theorem 2 in this form, with \(H_2(p) + D_2(p \,\|\, q)\) in place of \(H_2(p, q)\).

Remark 4 (Why models are fit by minimizing cross-entropy). Fitting a classifier by minimizing cross-entropy chooses the model’s distribution \(q\) so that a code built for \(q\) would spend as few bits as possible on the observed labels. By Remark 3, that pushes \(q\) toward the distribution the labels came from.

For a single labeled example, let \(e\) be the distribution that puts all its mass on the observed label. Then \(H_2(e) = 0\) (as for a coin that always lands heads), so the example’s cross-entropy loss \(H_2(e, q)\) equals \(D_2(e \,\|\, q)\): every bit the loss reports is a wasted one.

References

Cover, Thomas M., and Joy A. Thomas. 2006. Elements of Information Theory. 2nd ed. Wiley-Interscience. https://doi.org/10.1002/047174882X.

Shannon, C. E. 1948. “A Mathematical Theory of Communication.” Bell System Technical Journal 27 (3): 379–423. https://doi.org/10.1002/j.1538-7305.1948.tb01338.x.