Csr2011 june15 09_30_shen

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CSR-2011

Kolmogorov complexity as a language

Alexander ShenLIF CNRS, Marseille; on leave from

ИППИ РАН, Москва

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Kolmogorov complexity

I A powerful toolI Just a way to reformulate argumentsI three languages: combinatorial/algorithmic/probabilistic

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I A powerful tool

I Just a way to reformulate argumentsI three languages: combinatorial/algorithmic/probabilistic

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I A powerful toolI Just a way to reformulate arguments

I three languages: combinatorial/algorithmic/probabilistic

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I A powerful toolI Just a way to reformulate argumentsI three languages: combinatorial/algorithmic/probabilistic

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Reminder and notation

I K(x) = minimal length of a program that produces xI KD(x) = min|p| : D(p) = xI depends on the interpreter DI optimal D makes it minimal up to O(1) additive termI Variations:

p = string p = prefix of a sequencex = string plain prefix

K(x), C(x) KP(x), K(x)x = prefix decision monotone

of a sequence KR(x), KD(x) KM(x), Km(x)

Conditional complexity C(x|y):minimal length of a program p : y 7→ x.There is also a priori probability (in two versions: discrete, on strings;continuous, on prefixes)

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Reminder and notationI K(x) = minimal length of a program that produces x

I KD(x) = min|p| : D(p) = xI depends on the interpreter DI optimal D makes it minimal up to O(1) additive termI Variations:





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Reminder and notationI K(x) = minimal length of a program that produces xI KD(x) = min|p| : D(p) = x

I depends on the interpreter DI optimal D makes it minimal up to O(1) additive termI Variations:





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Reminder and notationI K(x) = minimal length of a program that produces xI KD(x) = min|p| : D(p) = xI depends on the interpreter D

I optimal D makes it minimal up to O(1) additive termI Variations:





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Reminder and notationI K(x) = minimal length of a program that produces xI KD(x) = min|p| : D(p) = xI depends on the interpreter DI optimal D makes it minimal up to O(1) additive term

I Variations:





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Reminder and notationI K(x) = minimal length of a program that produces xI KD(x) = min|p| : D(p) = xI depends on the interpreter DI optimal D makes it minimal up to O(1) additive termI Variations:





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Conditional complexity C(x|y):minimal length of a program p : y 7→ x.

There is also a priori probability (in two versions: discrete, on strings;continuous, on prefixes)

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Foundations of probability theory

I Random object or random process?I “well shuffled deck of cards”: any meaning?

[xkcd cartoon]

I randomness = incompressibility (maximal complexity)I ω = ω1ω2 . . . is random iff KM(ω1 . . . ωn) ≥ n−O(1)

I Classical probability theory: random sequence satisfies theStrong Law of Large Numbers with probability 1

I Algorithmic version: every (algorithmically) randomsequence satisfies SLLN

I algorithmic ⇒ classical: Martin-Löf random sequencesform a set of measure 1.

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I Random object or random process?

I “well shuffled deck of cards”: any meaning?

[xkcd cartoon]





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[xkcd cartoon]





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[xkcd cartoon]

I randomness = incompressibility (maximal complexity)

I ω = ω1ω2 . . . is random iff KM(ω1 . . . ωn) ≥ n−O(1)




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[xkcd cartoon]





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Sampling random strings (S.Aaronson)

I A device that (being switched on) produces N-bit string andstops

I “The device produces a random string”: what does itmean?

I classical: the output distribution is close to the uniform oneI effective: with high probability the output string isincompressible

I not equivalent if no assumptions about the deviceI but are related under some assumptions

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I classical: the output distribution is close to the uniform one

I effective: with high probability the output string isincompressible


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I not equivalent if no assumptions about the device

I but are related under some assumptions

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Example: matrices without uniform minors

I k× k minor of n× n Boolean matrix: select k rows and kcolumns

I minor is uniform if it is all-0 or all-1.I claim: there is a n× n bit matrix without k× k uniformminors for k = 3 log n.

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Counting argument and complexity reformulation

I ≤ nk × nk positions of the minor [k = 3 logn]I 2 types of uniform minors (0/1)I 2n

2−k2 possibilities for the restI n2k × 2× 2n

2−k2 = 2log n×2×3 log n+1+(n2−9 log2 n) < 2n2 .

I log n bits to specify a column or row: 2k log n bits in totalI one additional bit to specify the type of minor (0/1)I n2 − k2 bits to specify the rest of the matrixI 2k log n+ 1+ (n2 − k2) = 6 log2 n+ 1+ (n2 − 9 log2 n) < n2.

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I ≤ nk × nk positions of the minor [k = 3 logn]

I 2 types of uniform minors (0/1)I 2n




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I ≤ nk × nk positions of the minor [k = 3 logn]I 2 types of uniform minors (0/1)

I 2n2−k2 possibilities for the rest

I n2k × 2× 2n2−k2 = 2log n×2×3 log n+1+(n2−9 log2 n) < 2n

2 .


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2−k2 possibilities for the rest

I n2k × 2× 2n2−k2 = 2log n×2×3 log n+1+(n2−9 log2 n) < 2n

2 .


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2−k2 =

2log n×2×3 log n+1+(n2−9 log2 n) < 2n2 .


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I log n bits to specify a column or row: 2k log n bits in total

I one additional bit to specify the type of minor (0/1)I n2 − k2 bits to specify the rest of the matrixI 2k log n+ 1+ (n2 − k2) = 6 log2 n+ 1+ (n2 − 9 log2 n) < n2.

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I log n bits to specify a column or row: 2k log n bits in totalI one additional bit to specify the type of minor (0/1)

I n2 − k2 bits to specify the rest of the matrixI 2k log n+ 1+ (n2 − k2) = 6 log2 n+ 1+ (n2 − 9 log2 n) < n2.

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I log n bits to specify a column or row: 2k log n bits in totalI one additional bit to specify the type of minor (0/1)I n2 − k2 bits to specify the rest of the matrix

I 2k log n+ 1+ (n2 − k2) = 6 log2 n+ 1+ (n2 − 9 log2 n) < n2.

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One-tape Turing machines

I copying n-bit string on 1-tape TM requires Ω(n2) time

I complexity version: if initially the tape was empty on theright of the border, then after n steps the complexity of azone that is d cells far from the border is O(n/d).

K(u(t)) ≤ O(n/d)I proof: border guards in each cell of the border security zone write down the

contents of the head of TM; each of the records is enough to reconstruct u(t) so

the length of it should be Ω(K(u(t)); the sum of lengths does not exceed time

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One-tape Turing machinesI copying n-bit string on 1-tape TM requires Ω(n2) time





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K(u(t)) ≤ O(n/d)

I proof: border guards in each cell of the border security zone write down the



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Everywhere complex sequences

I Random sequence has n-bit prefix of complexity nI but some factors (substrings) have small complexityI Levin: there exist everywhere complex sequences: everyn-bit substring has complexity 0.99n−O(1)

I Combinatorial equivalent: Let F be a set of strings that hasat most 20.99n strings of length n. Then there is a sequenceω s.t. all sufficiently long substrings of ω are not in F.

I combinatorial and complexity proofs not just translations ofeach other (Lovasz lemma, Rumyantsev, Miller, Muchnik)

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I Random sequence has n-bit prefix of complexity n

I but some factors (substrings) have small complexityI Levin: there exist everywhere complex sequences: everyn-bit substring has complexity 0.99n−O(1)



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I Random sequence has n-bit prefix of complexity nI but some factors (substrings) have small complexity

I Levin: there exist everywhere complex sequences: everyn-bit substring has complexity 0.99n−O(1)



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I Random sequence has n-bit prefix of complexity nI but some factors (substrings) have small complexityI Levin: there exist everywhere complex sequences: everyn-bit substring has complexity 0.99n−O(1)



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Gilbert-Varshamov complexity bound

I coding theory: how many n-bit strings x1, . . . , xk one canfind if Hamming distance between every two is at least d

I lower bound (Gilbert–Varshamov)I then < d/2 changed bits are harmlessI but bit insertion or deletions could beI general requirement: C(xi|xj) ≥ dI generalization of GV bound: d-separated family of size

Ω(2n−d)

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I lower bound (Gilbert–Varshamov)

I then < d/2 changed bits are harmlessI but bit insertion or deletions could beI general requirement: C(xi|xj) ≥ dI generalization of GV bound: d-separated family of size

Ω(2n−d)

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I lower bound (Gilbert–Varshamov)I then < d/2 changed bits are harmless

I but bit insertion or deletions could beI general requirement: C(xi|xj) ≥ dI generalization of GV bound: d-separated family of size

Ω(2n−d)

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I lower bound (Gilbert–Varshamov)I then < d/2 changed bits are harmlessI but bit insertion or deletions could be

I general requirement: C(xi|xj) ≥ dI generalization of GV bound: d-separated family of size

Ω(2n−d)

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I lower bound (Gilbert–Varshamov)I then < d/2 changed bits are harmlessI but bit insertion or deletions could beI general requirement: C(xi|xj) ≥ d

I generalization of GV bound: d-separated family of sizeΩ(2n−d)

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I lower bound (Gilbert–Varshamov)I then < d/2 changed bits are harmlessI but bit insertion or deletions could beI general requirement: C(xi|xj) ≥ dI generalization of GV bound: d-separated family of size

Ω(2n−d)

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Inequalities for complexities and combinatorialinterpretation

I C(x, y) ≤ C(x) + C(y|x) +O(log)

I C(x, y) ≥ C(x) + C(y|x) +O(log)I C(x, y) < k+ l ⇒ C(x) < k+O(log) or C(y|x) < l+O(log)I every set A of size < 2k+l can be split into two partsA = A1 ∪ A2 such that w(A1) ≤ 2k and h(A2) ≤ 2l

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I C(x, y) ≤ C(x) + C(y|x) +O(log)

I C(x, y) ≥ C(x) + C(y|x) +O(log)

I C(x, y) < k+ l ⇒ C(x) < k+O(log) or C(y|x) < l+O(log)I every set A of size < 2k+l can be split into two partsA = A1 ∪ A2 such that w(A1) ≤ 2k and h(A2) ≤ 2l

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I C(x, y) ≤ C(x) + C(y|x) +O(log)

I C(x, y) ≥ C(x) + C(y|x) +O(log)I C(x, y) < k+ l ⇒ C(x) < k+O(log) or C(y|x) < l+O(log)

I every set A of size < 2k+l can be split into two partsA = A1 ∪ A2 such that w(A1) ≤ 2k and h(A2) ≤ 2l

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I C(x, y) ≤ C(x) + C(y|x) +O(log)

I C(x, y) ≥ C(x) + C(y|x) +O(log)I C(x, y) < k+ l ⇒ C(x) < k+O(log) or C(y|x) < l+O(log)I every set A of size < 2k+l can be split into two partsA = A1 ∪ A2 such that w(A1) ≤ 2k and h(A2) ≤ 2l

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One more inequality

I 2C(x, y, z) ≤ C(x, y) + C(y, z) + C(x, z)

I V2 ≤ S1 × S2 × S3

I Also for Shannon entropies; special case of Shearerlemma

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Common information and graph minors

I mutual information: I(a : b) = C(a) + C(b)− C(a,b)

I common information:

I combinatorial: graph minors

can the graph be covered by 2δ minors of size 2α−δ × 2β−δ?

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Almost uniform sets

I nonuniformity= (maximal section)/(average section)

I Theorem: every set of N elements can be represented asunion of polylog(N) sets whose nonuniformity is polylog(N).

I multidimensional versionI how to construct parts using Kolmogorov complexity: takestrings with given complexity bounds

I so simple that it is not clear what is the combinatorialtranslation

I but combinatorial argument exists (and gives even astronger result)

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Almost uniform setsI nonuniformity= (maximal section)/(average section)





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I multidimensional version

I how to construct parts using Kolmogorov complexity: takestrings with given complexity bounds



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Shannon coding theorem

I ξ is a random variable; k values, probabilities p1, . . . , pkI ξN: N independent trials of ξI Shannon’s informal question: how many bits are needed toencode a “typical” value of ξN?

I Shannon’s answer: NH(ξ), where

H(ξ) = p1 log(1/p1) + . . .+ pn log(1/pn).

I formal statement is a bit complicatedI Complexity version: with high probablity the value of ξNhas complexity close to NH(ξ).

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I ξ is a random variable; k values, probabilities p1, . . . , pk

I ξN: N independent trials of ξI Shannon’s informal question: how many bits are needed toencode a “typical” value of ξN?


H(ξ) = p1 log(1/p1) + . . .+ pn log(1/pn).


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I ξ is a random variable; k values, probabilities p1, . . . , pkI ξN: N independent trials of ξ

I Shannon’s informal question: how many bits are needed toencode a “typical” value of ξN?


H(ξ) = p1 log(1/p1) + . . .+ pn log(1/pn).


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H(ξ) = p1 log(1/p1) + . . .+ pn log(1/pn).


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H(ξ) = p1 log(1/p1) + . . .+ pn log(1/pn).

I formal statement is a bit complicated

I Complexity version: with high probablity the value of ξNhas complexity close to NH(ξ).

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H(ξ) = p1 log(1/p1) + . . .+ pn log(1/pn).


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Complexity, entropy and group size

I 2C(x, y, z) ≤ C(x, y) + C(y, z) + C(x, z) +O(log)I The same for entropy:

2H(ξ, η, τ) ≤ H(ξ, η) + H(ξ, τ) + H(η, τ)I …and even for the sizes of subgroups U,V,W of somefinite group G: 2 log(|G|/|U ∩ V ∩W|) ≤log(|G|/|U ∩ V|) + log(|G|/|U ∩W|) + log(|G|/|V ∩W|).

I in all three cases inequalities are the same(Romashchenko, Chan, Yeung)

I some of them are quite strange:

I(a : b) ≤≤ I(a : b|c)+I(a : b|d)+I(c : d)+I(a : b|e)+I(a : e|b)+I(b : e|a)

I Related to Romashchenko’s theorem: if three last termsare zeros, one can extract common information from a,b,e.

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I 2C(x, y, z) ≤ C(x, y) + C(y, z) + C(x, z) +O(log)

I The same for entropy:2H(ξ, η, τ) ≤ H(ξ, η) + H(ξ, τ) + H(η, τ)

I …and even for the sizes of subgroups U,V,W of somefinite group G: 2 log(|G|/|U ∩ V ∩W|) ≤log(|G|/|U ∩ V|) + log(|G|/|U ∩W|) + log(|G|/|V ∩W|).





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2H(ξ, η, τ) ≤ H(ξ, η) + H(ξ, τ) + H(η, τ)

I …and even for the sizes of subgroups U,V,W of somefinite group G: 2 log(|G|/|U ∩ V ∩W|) ≤log(|G|/|U ∩ V|) + log(|G|/|U ∩W|) + log(|G|/|V ∩W|).





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2H(ξ, η, τ) ≤ H(ξ, η) + H(ξ, τ) + H(η, τ)I …and even for the sizes of subgroups U,V,W of somefinite group G: 2 log(|G|/|U ∩ V ∩W|) ≤log(|G|/|U ∩ V|) + log(|G|/|U ∩W|) + log(|G|/|V ∩W|).





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Muchnik and Slepian–Wolf

I a,b: two stringsI we look for a program p that maps a to bI by definition C(p) is at least C(b|a) but could be higherI there exist p : a 7→ b that is simple relative to b, e.g., “mapeverything to b”

I Muchnik theorem: it is possible to combine these twoconditions: there exists p : a 7→ b such that C(p) ≈ C(b|a)and C(p|b) ≈ 0

I information theory analog: Wolf–SlepianI similar technique was developed by Fortnow and Laplante(randomness extractors)

I (Romashchenko, Musatov): how to use explicit extractorsand derandomization to get space-bounded versions

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I a,b: two strings

I we look for a program p that maps a to bI by definition C(p) is at least C(b|a) but could be higherI there exist p : a 7→ b that is simple relative to b, e.g., “mapeverything to b”




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I a,b: two stringsI we look for a program p that maps a to b

I by definition C(p) is at least C(b|a) but could be higherI there exist p : a 7→ b that is simple relative to b, e.g., “mapeverything to b”




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I a,b: two stringsI we look for a program p that maps a to bI by definition C(p) is at least C(b|a) but could be higher

I there exist p : a 7→ b that is simple relative to b, e.g., “mapeverything to b”




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I information theory analog: Wolf–Slepian

I similar technique was developed by Fortnow and Laplante(randomness extractors)


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Computability theory: simple sets

I Simple set: enumerable set with infinite complement, butno algorithm can generate infinitely many elements fromthe complement

I Construction using Kolmogorov complexity: a simple stringx has C(x) ≤ |x|/2.

I Most strings are not simple ⇒ infinite complementI Let x1, x2, . . . be a computable sequence of differentnon-simple strings

I May assume wlog that |xi| > i and therefore C(xi) > i/2I but to specify xi we need O(log i) bits onlyI “Minimal integer that cannot be described in ten Englishwords” (Berry)

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I Most strings are not simple ⇒ infinite complement

I Let x1, x2, . . . be a computable sequence of differentnon-simple strings


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I May assume wlog that |xi| > i and therefore C(xi) > i/2

I but to specify xi we need O(log i) bits onlyI “Minimal integer that cannot be described in ten Englishwords” (Berry)

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I May assume wlog that |xi| > i and therefore C(xi) > i/2I but to specify xi we need O(log i) bits only

I “Minimal integer that cannot be described in ten Englishwords” (Berry)

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Lower semicomputable random reals

I∑

ai computable converging series with rational termsI is α =

∑ai computable (ε → ε-approximation)?

I not necessarily (Specker example)I lower semicomputable realsI Solovay classification: α ≼ β if ε-approximation to β can beeffectively converted to O(ε)-approximation to α

I There are maximal elements = random lowersemicomputable reals

I = slowly converging seriesI modulus of convergence: ε 7→ N(ε) = how many terms areneeded for ε-precision

I maximal elements: N(2−n) > BP(n−O(1)), where BP(k)is the maximal integer whose prefix complexity is k or less.

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I∑

ai computable converging series with rational terms

I is α =∑

ai computable (ε → ε-approximation)?I not necessarily (Specker example)I lower semicomputable realsI Solovay classification: α ≼ β if ε-approximation to β can beeffectively converted to O(ε)-approximation to α




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I∑







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I∑



I not necessarily (Specker example)

I lower semicomputable realsI Solovay classification: α ≼ β if ε-approximation to β can beeffectively converted to O(ε)-approximation to α




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I∑



I not necessarily (Specker example)I lower semicomputable reals

I Solovay classification: α ≼ β if ε-approximation to β can beeffectively converted to O(ε)-approximation to α




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I∑







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I∑




I There are maximal elements

= random lowersemicomputable reals



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I∑







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I∑





I = slowly converging series

I modulus of convergence: ε 7→ N(ε) = how many terms areneeded for ε-precision


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I∑







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Lovasz local lemma: constructive proof

I CNF: (a ∧ ¬b ∧ . . .) ∨ (¬c ∧ e ∧ . . .) ∨ . . .

I neighbors: clauses having common variablesI several clauses with k literals in eachI each clause has o(2k) neighborsI ⇒ CNF is satisfiableI Non-constructive proof: lower bound for probability, Lovaszlocal lemma

I Naïve algorithm: just resample false clause (while theyexist)

I Recent breakthrough (Moser): this algorithm with highprobability terminates quickly

I Explanation: if not, the sequence of resampled clauseswould encode the random bits used in resampling makingthem compressible

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I CNF: (a ∧ ¬b ∧ . . .) ∨ (¬c ∧ e ∧ . . .) ∨ . . .

I neighbors: clauses having common variables

I several clauses with k literals in eachI each clause has o(2k) neighborsI ⇒ CNF is satisfiableI Non-constructive proof: lower bound for probability, Lovaszlocal lemma




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I CNF: (a ∧ ¬b ∧ . . .) ∨ (¬c ∧ e ∧ . . .) ∨ . . .

I neighbors: clauses having common variablesI several clauses with k literals in each

I each clause has o(2k) neighborsI ⇒ CNF is satisfiableI Non-constructive proof: lower bound for probability, Lovaszlocal lemma




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I CNF: (a ∧ ¬b ∧ . . .) ∨ (¬c ∧ e ∧ . . .) ∨ . . .

I neighbors: clauses having common variablesI several clauses with k literals in eachI each clause has o(2k) neighbors

I ⇒ CNF is satisfiableI Non-constructive proof: lower bound for probability, Lovaszlocal lemma




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I CNF: (a ∧ ¬b ∧ . . .) ∨ (¬c ∧ e ∧ . . .) ∨ . . .

I neighbors: clauses having common variablesI several clauses with k literals in eachI each clause has o(2k) neighborsI ⇒ CNF is satisfiable

I Non-constructive proof: lower bound for probability, Lovaszlocal lemma




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I CNF: (a ∧ ¬b ∧ . . .) ∨ (¬c ∧ e ∧ . . .) ∨ . . .

I neighbors: clauses having common variablesI several clauses with k literals in eachI each clause has o(2k) neighborsI ⇒ CNF is satisfiableI Non-constructive proof: lower bound for probability, Lovaszlocal lemma




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Berry, Gödel, Chaitin, Raz

I There are only finitely many strings of complexity < nI for all strings (except finitely many ones) x the statementK(x) > n is true

I Can all true statements of this form be provable?I No, otherwise we could effectively generate string ofcomplexity > n by enumerating all proofs

I and get Berry’s paradox: the first provable statementC(x) > n for given n gives some x of complexity > n thatcan be described by O(log n) bits

I (Gödel theorem in Chaitin form): There are only finitelymany n such that C(x) > n is provable for some x. (Notethat ∃xC(x) > n is always provable!)

I (Gödel second theorem, Kritchman–Raz proof): the“unexpected test paradox” (a test will be given next weekbut it won’t be known before the day of the test)

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Berry, Gödel, Chaitin, RazI There are only finitely many strings of complexity < n

I for all strings (except finitely many ones) x the statementK(x) > n is true





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Berry, Gödel, Chaitin, RazI There are only finitely many strings of complexity < nI for all strings (except finitely many ones) x the statementK(x) > n is true





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I Can all true statements of this form be provable?

I No, otherwise we could effectively generate string ofcomplexity > n by enumerating all proofs




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I (Gödel theorem in Chaitin form): There are only finitelymany n such that C(x) > n is provable for some x.

(Notethat ∃xC(x) > n is always provable!)


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13th Hilbert’s problem

I Function of ≥ 3 variables: e.g., solution of a polynomial ofdegree 7 as function of its coefficients

I Is it possible to represent this function as a composition offunctions of at most 2 variables?

I Yes, with weird functions (Cantor)I Yes, even with continuous functions (Kolmogorov, Arnold)I Circuit version: explicit function Bn × Bn × Bn → Bn

(polynomial circuit size?) that cannot be represented ascomposition of O(1) functions Bn × Bn → Bn. Not known.

I Kolmogorov complexity version: we have three stringsa,b, c on a blackboard. It is allowed to write (add) a newstring if it simple relative to two of the strings on the board.Which strings we can obtain in O(1) steps? Only strings ofsmall complexity relative to a,b, c, but not all of them (forrandom a,b, c)

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I Yes, with weird functions (Cantor)

I Yes, even with continuous functions (Kolmogorov, Arnold)I Circuit version: explicit function Bn × Bn × Bn → Bn



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I Yes, with weird functions (Cantor)I Yes, even with continuous functions (Kolmogorov, Arnold)

I Circuit version: explicit function Bn × Bn × Bn → Bn



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I Yes, with weird functions (Cantor)I Yes, even with continuous functions (Kolmogorov, Arnold)I Circuit version: explicit function Bn × Bn × Bn → Bn



. . . . . .

Secret sharing

I secret s and three people; any two are able to reconstructthe secret working together; but each one in isolation hasno information about it

I assume s ∈ F (a field); take random a and tell the peoplea, a+ s and a+ 2s (assuming 2 = 0 in F)

I other secret sharing schemes; how long should be secrets(not understood)

I Kolmogorov complexity settings: for a given s find a,b, csuch that

C(s|a),C(s|b),C(s|c) ≈ C(s); C(s|a,b),C(s|a, c),C(s|b, c) ≈ 0

I Some relation between Kolmogorov and traditional settings(Romashchenko, Kaced)

. . . . . .

Quasi-cryptography

I Alice has some information aI Bob wants to let her know some bI by sending some message fI in such a way that Eve gets minimal information about bI Formally: for given a, b find f such that C(b|a, f) ≈ 0 andC(b|f) → max.

I Theorem (Muchnik): it is always possible to haveC(b|f) ≈ min(C(b),C(a))

I Full version: Eve knows some c and we want to sendmessage of a minimal possible length C(b|a)

. . . . . .

Quasi-cryptography

I Alice has some information a

I Bob wants to let her know some bI by sending some message fI in such a way that Eve gets minimal information about bI Formally: for given a, b find f such that C(b|a, f) ≈ 0 andC(b|f) → max.



. . . . . .

Quasi-cryptography

I Alice has some information aI Bob wants to let her know some b

I by sending some message fI in such a way that Eve gets minimal information about bI Formally: for given a, b find f such that C(b|a, f) ≈ 0 andC(b|f) → max.



. . . . . .

Quasi-cryptography

I Alice has some information aI Bob wants to let her know some bI by sending some message f

I in such a way that Eve gets minimal information about bI Formally: for given a, b find f such that C(b|a, f) ≈ 0 andC(b|f) → max.



. . . . . .

Quasi-cryptography

I Alice has some information aI Bob wants to let her know some bI by sending some message fI in such a way that Eve gets minimal information about b

I Formally: for given a, b find f such that C(b|a, f) ≈ 0 andC(b|f) → max.



. . . . . .

Quasi-cryptography

I Alice has some information aI Bob wants to let her know some bI by sending some message fI in such a way that Eve gets minimal information about bI Formally: for given a, b find f such that C(b|a, f) ≈ 0 andC(b|f) → max.



. . . . . .

Andrej Muchnik (1958-2007)

Csr2011 june15 09_30_shen

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