5: Pseudorandom Generators
- Page ID
- 7347
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\(\newcommand{\avec}{\mathbf a}\) \(\newcommand{\bvec}{\mathbf b}\) \(\newcommand{\cvec}{\mathbf c}\) \(\newcommand{\dvec}{\mathbf d}\) \(\newcommand{\dtil}{\widetilde{\mathbf d}}\) \(\newcommand{\evec}{\mathbf e}\) \(\newcommand{\fvec}{\mathbf f}\) \(\newcommand{\nvec}{\mathbf n}\) \(\newcommand{\pvec}{\mathbf p}\) \(\newcommand{\qvec}{\mathbf q}\) \(\newcommand{\svec}{\mathbf s}\) \(\newcommand{\tvec}{\mathbf t}\) \(\newcommand{\uvec}{\mathbf u}\) \(\newcommand{\vvec}{\mathbf v}\) \(\newcommand{\wvec}{\mathbf w}\) \(\newcommand{\xvec}{\mathbf x}\) \(\newcommand{\yvec}{\mathbf y}\) \(\newcommand{\zvec}{\mathbf z}\) \(\newcommand{\rvec}{\mathbf r}\) \(\newcommand{\mvec}{\mathbf m}\) \(\newcommand{\zerovec}{\mathbf 0}\) \(\newcommand{\onevec}{\mathbf 1}\) \(\newcommand{\real}{\mathbb R}\) \(\newcommand{\twovec}[2]{\left[\begin{array}{r}#1 \\ #2 \end{array}\right]}\) \(\newcommand{\ctwovec}[2]{\left[\begin{array}{c}#1 \\ #2 \end{array}\right]}\) \(\newcommand{\threevec}[3]{\left[\begin{array}{r}#1 \\ #2 \\ #3 \end{array}\right]}\) \(\newcommand{\cthreevec}[3]{\left[\begin{array}{c}#1 \\ #2 \\ #3 \end{array}\right]}\) \(\newcommand{\fourvec}[4]{\left[\begin{array}{r}#1 \\ #2 \\ #3 \\ #4 \end{array}\right]}\) \(\newcommand{\cfourvec}[4]{\left[\begin{array}{c}#1 \\ #2 \\ #3 \\ #4 \end{array}\right]}\) \(\newcommand{\fivevec}[5]{\left[\begin{array}{r}#1 \\ #2 \\ #3 \\ #4 \\ #5 \\ \end{array}\right]}\) \(\newcommand{\cfivevec}[5]{\left[\begin{array}{c}#1 \\ #2 \\ #3 \\ #4 \\ #5 \\ \end{array}\right]}\) \(\newcommand{\mattwo}[4]{\left[\begin{array}{rr}#1 \amp #2 \\ #3 \amp #4 \\ \end{array}\right]}\) \(\newcommand{\laspan}[1]{\text{Span}\{#1\}}\) \(\newcommand{\bcal}{\cal B}\) \(\newcommand{\ccal}{\cal C}\) \(\newcommand{\scal}{\cal S}\) \(\newcommand{\wcal}{\cal W}\) \(\newcommand{\ecal}{\cal E}\) \(\newcommand{\coords}[2]{\left\{#1\right\}_{#2}}\) \(\newcommand{\gray}[1]{\color{gray}{#1}}\) \(\newcommand{\lgray}[1]{\color{lightgray}{#1}}\) \(\newcommand{\rank}{\operatorname{rank}}\) \(\newcommand{\row}{\text{Row}}\) \(\newcommand{\col}{\text{Col}}\) \(\renewcommand{\row}{\text{Row}}\) \(\newcommand{\nul}{\text{Nul}}\) \(\newcommand{\var}{\text{Var}}\) \(\newcommand{\corr}{\text{corr}}\) \(\newcommand{\len}[1]{\left|#1\right|}\) \(\newcommand{\bbar}{\overline{\bvec}}\) \(\newcommand{\bhat}{\widehat{\bvec}}\) \(\newcommand{\bperp}{\bvec^\perp}\) \(\newcommand{\xhat}{\widehat{\xvec}}\) \(\newcommand{\vhat}{\widehat{\vvec}}\) \(\newcommand{\uhat}{\widehat{\uvec}}\) \(\newcommand{\what}{\widehat{\wvec}}\) \(\newcommand{\Sighat}{\widehat{\Sigma}}\) \(\newcommand{\lt}{<}\) \(\newcommand{\gt}{>}\) \(\newcommand{\amp}{&}\) \(\definecolor{fillinmathshade}{gray}{0.9}\)One-time pad requires a key that’s as long as the plaintext. Let’s forget that we know about this limitation. Suppose Alice & Bob share only a short \(\lambda\)-bit secret \(k\), but they want to encrypt a \(2 \lambda\)-bit plaintext \(m\). They don’t know that (perfect) one-time secrecy is impossible in this setting (Exercise 2.11), so they try to get it to work anyway using the following reasoning:
- The only encryption scheme they know about is one-time pad, so they decide that the ciphertext will have the form \(c=m \,\oplus\) ??. This means that the unknown value \(? ?\) must be \(2 \lambda\) bits long.
- In order for the security of one-time pad to apply, the unknown value ?? should be uniformly distributed.
- The process of obtaining the unknown value ?? from the shared key \(k\) should be deterministic, so that the sender and receiver compute the same value and decryption works correctly.
Let \(G\) denote the process that transforms the key \(k\) into this mystery value. Then \(G\) : \(\{0,1\}^{\lambda} \rightarrow\{0,1\}^{2 \lambda}\), and the encryption scheme is \(\operatorname{Enc}(k, m)=m \oplus G(k)\)
It is not hard to see that if \(G\) is a deterministic function, then there are only \(2^{\lambda}\) possible outputs of \(G\), so the distribution of \(G(k)\) cannot be uniform in \(\{0,1\}^{2 \lambda}\). We therefore cannot argue that the scheme is secure in the same way as one-time pad.
However, what if the distribution of \(G(k)\) values is not perfectly uniform but only "close enough" to uniform? Suppose no polynomial-time algorithm can distinguish the distribution of \(G(k)\) values from the uniform distribution. Then surely this ought to be "close enough" to uniform for practical purposes. This is exactly the idea of pseudorandomness. It turns out that if \(G\) has a pseudorandomness property, then the encryption scheme described above is actually secure (against polynomial-time adversaries, in the sense discussed in the previous chapter).