communication-rounds tradeoffs for common …communication-rounds tradeo s for common randomness and...

Communication-Rounds Tradeoffs forCommon Randomness and Secret Key

Generation

Mitali Bafna†, Badih Ghazi∗,Noah Golowich†, and Madhu Sudan†

† Harvard University ∗ Google Research

January 8, 2019

1

Common Randomness Generation (CRG)

(Xi ,Yi) ∼ µ (source), Alice ← Xi , Bob ← Yi .

Several rounds of commun., starting w/ Alice:I m1 = m1(Xi);I m2 = m2(m1, Yi);I m3 = m3(Xi,m1,m2), . . ..

At end: Alice & Bob output KA,KB , resp., s.t.:I KA = KA(Xi,m1, . . . ,mr );

KB = KB(Yi,m1, . . . ,mr ).I KA = KB w.h.p., andI KA,KB have large min-entropy.

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(Xi ,Yi) ∼ µ (source), Alice ← Xi , Bob ← Yi .Several rounds of commun., starting w/ Alice:

I m1 = m1(Xi);I m2 = m2(m1, Yi);I m3 = m3(Xi,m1,m2), . . ..



2





KB = KB(Yi,m1, . . . ,mr ).

I KA = KB w.h.p., andI KA,KB have large min-entropy.

2






2

Secret Key Generation (SKG)

Alice Bob

Eve

Y1, Y2, Y3 ...

X1, X2, X3 ...

Same as CRG, but key must be secure againsteavesdropper Eve that watches the communication.

Question: Are there sources µ such thathaving more rounds can lead to more efficient(i.e. lower communication) protocols for CRGor SKG?

3

Background: 1-Round & 2-Round Communication

[Ahlswede & Csiszar, ’93 & ’98]: CRG and SKG for1-round, 2-round protocols in amortized setting:For ε→ 0, characterize (H ,C ) pairs s.t.

I Samples: (XN ,Y N) ∼ µ⊗N ;I Communication: (C + ε) · N ;I Entropy of key: (H − ε) · N .

[Guruswami & Radhakrishnan, ’16] & [Ghazi & Jayram, ’18]:non-amortized setting (our work): near-optimaltradeoff between communication, key length, &agreement probability for “simple” sources: BSC,BEC, BGS.

4

Background: Multi-Round Communication

Generalization of Ahlswede & Csiszarcharacterization to multi-round protocols by [Tyagi,

’13] & [Liu et al., ’16].

For binary channels, additional rounds does nothelp reduce communication.

[Tyagi, ’13]: ternary source s.t. 2-round protocolsrequire less communication than 1-round protocols.

Otherwise: no rounds vs. communication tradeoffs(incl. in non-amortized case)!

5

Formalizing Non-Amortized Setting

Definition ((r , c)-protocol)Π is (r , c)-protocol if:

≤ r rounds (messages).

≤ c bits total.

L represents number of bits of entropy in random keys:

Definition (L-CRG)Π gives L-CRG if Alice, Bob, given samples(X1,Y1), . . . , (XT ,YT ) ∼ µ⊗T , output KA,KB , s.t.:

minH∞(KA),H∞(KB) ≥ L.

Pr[KA = KB ] ≥ 1− ε.

6

Formalizing Non-Amortized Setting

Definition ((r , c)-protocol)Π is (r , c)-protocol if:

≤ r rounds (messages).

≤ c bits total.

L represents number of bits of entropy in random keys:

Definition (L-CRG)Π gives L-CRG if Alice, Bob, given samples(X1,Y1), . . . , (XT ,YT ) ∼ µ⊗T , output KA,KB , s.t.:

minH∞(KA),H∞(KB) ≥ L.

Pr[KA = KB ] ≥ 1− ε.6

r vs. r/2 gap

Question: ∀δ > 0, r , L, ∃µ s.t.

∃(r , δL)-protocol for L-CRG from µ;

No (r − 1, (1− δ)L)-protocol for L-CRG from µ?

Theorem (BGGS, ’19)∀n, r , L construct µr ,n,L s.t.

∃(r + 1,O(r log n))-protocol for L-CRG from µr ,n,L;

No(

r2 − 2,min

o(L), n

poly log(n)

)-protocol for

L-CRG from µr ,n,L.

Also: same theorem for L-SKG!

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Pointer-chasing source for CRG: µr ,n,L

i0 ~ [n]

π1 , .., π

r ~ S

n

A1 , …, A

n ~ 0,1L

B1 , …, B

n ~ 0,1L

Bob's inputs = (i

0, π

2 , π

4 , …, B

1 , …, B

n )

Alice's inputs = (π

1 , π

3 , …, A

1 , …, A

n )

8


π1

i0 ~ [n]

π1 , .., π

r ~ S

n

A1 , …, A

n ~ 0,1L

B1 , …, B

n ~ 0,1L

ij+1

= πj+1

(ij).

ij = “pointers”

i0

π2

π3

πr-1

πr

Bob's inputs = (i

0, π

2 , π

4 , …, B

1 , …, B

n )


1 , π

3 , …, A

1 , …, A

n )

8


π1

i0 ~ [n]

π1 , .., π

r ~ S

n

A1 , …, A

n ~ 0,1L

B1 , …, B

n ~ 0,1L

Where Ai = B

i ,

ij+1

= πj+1

(ij).

ij = “pointers”

r r

i0

π2

π3

πr-1

πr

An

A1

A3...A

2B

nB

3...B

2B

1

Bob's inputs = (i

0, π

2 , π

4 , …, B

1 , …, B

n )


1 , π

3 , …, A

1 , …, A

n )

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Background: Standard Pointer Chasing Problem

[Duris et al., ’84] & [Nisan & Wigderson, ’93], etc.:

π1

i0

π2

π3

πr-1

πr

Bob's inputs = (i

0, π

2 , π

4 , …, π

r-1 )


1 , π

3 , …, π

r )

ir = π

r ... π

1(i

0) =

?

i0 ~ [n]

π1 , .., π

r ~ S

n 9

Upper/lower bounds for CRG from µr ,n,L

Upper bound: follow pointers:(r + 1) rounds,communication (r + 1)dlog ne.

π1

i0

π2

π3

πr-1

πr

An

A1

A3...A

2B

nB

3...B

2B

1

Bob's inputs = (i, π

2 , π

4 , …, B

1 , …, B

n )


1 , π

3 , …, A

1 , …, A

n )

Lower bound for protocols w/ ≤ r rounds:I [Duris et al., ’84] & [Nisan & Wigderson, ’93], etc.:

Pointer chasing is hard (for det. & rand. protocols).I Problem: don’t have to solve pointer chasing for

L-CRG!I “Guess” index j such that Aj = Bj ⇒ dlog ne-bit

non-deterministic protocol!I Modular solution?

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π1

i0

π2

π3

πr-1

πr

An

A1

A3...A

2B

nB

3...B

2B

1


2 , π

4 , …, B

1 , …, B

n )


1 , π

3 , …, A

1 , …, A

n )


Pointer chasing is hard (for det. & rand. protocols).

I Problem: don’t have to solve pointer chasing forL-CRG!

I “Guess” index j such that Aj = Bj ⇒ dlog ne-bitnon-deterministic protocol!

I Modular solution?

10



π1

i0

π2

π3

πr-1

πr

An

A1

A3...A

2B

nB

3...B

2B

1


2 , π

4 , …, B

1 , …, B

n )


1 , π

3 , …, A

1 , …, A

n )




non-deterministic protocol!

I Modular solution?

10



π1

i0

π2

π3

πr-1

πr

An

A1

A3...A

2B

nB

3...B

2B

1


2 , π

4 , …, B

1 , …, B

n )


1 , π

3 , …, A

1 , …, A

n )




non-deterministic protocol!I Modular solution?

10

Reduction 1: Indistinguishability

Marginals of µr ,n,L: X = (π1, π3, . . . , πr ,A1, . . . ,An),Y = (i0, π2, . . . , πr−1,B1, . . . ,Bn).

Π distinguishes µ, ν if TVD(Πµ,Πν) ≥ const.

ClaimSuppose c L. If ∃ (r/2− 2, c) protocol for L-CRGfrom µ = µr ,n,L, then µ & µX × µY are distinguishable by(r/2− 1)-round protocol w/ communication c + O(1).

Idea of proof: L-CRG is impossible with communication L when parties have indep. inputs (e.g., [Cannone et al.,

’17]).

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Overview of Reductions

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Reduction 2: Pointer Verification

PVYES(r , n),PVNO(r , n) distributions on (X , Y ):X = (π1, π3, . . . , πr), Y = (i0, j0, π2, π4, . . . , πr−1):

PVYES: j0 = πr · · · π1(i0). PVNO: j0 random.

Protocol distinguishing µ = µr ,n,L & µX × µYgives protocol distinguishing PVYES & PVNO

(using Ω(n) disjointness lower bound).

13

Reduction 2: Pointer Verification

PVYES(r , n),PVNO(r , n) distributions on (X , Y ):X = (π1, π3, . . . , πr), Y = (i0, j0, π2, π4, . . . , πr−1):

PVYES: j0 = πr · · · π1(i0). PVNO: j0 random.

Protocol distinguishing µ = µr ,n,L & µX × µYgives protocol distinguishing PVYES & PVNO

(using Ω(n) disjointness lower bound).13

Overview of Reductions

14

Indistinguishability of PVYES, PVNO

r , n implicit: PVYES = PVYES(n, r), PVNO = PVNO(n, r):

Theorem (BGGS, ’19)

∀r , n, PVNO & PVYES are(r−12 ,

npoly log(n)

)-indisting.

Why only r−12 rounds? ∃

(r+12 ,O(r log n)

)-protocol:

PVNO: PVYES:

15





npoly log(n)

)-indisting.

Why only r−12 rounds?

∃(r+12 ,O(r log n)

)-protocol:

PVNO: PVYES:

15





npoly log(n)

)-indisting.

Why only r−12 rounds? ∃

(r+12 ,O(r log n)

)-protocol:

PVNO: PVYES:

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Proof of indist. of PVYES & PVNO

Idea: “round elimination” ([Nisan & Wigderson, ’93]).

Dealing with non-independence of inputs underPVYES makes proof more difficult.1

Idea: “peel” off 2 permutations (1 round) at a time,maintaining invariants:

I H(π1, . . . , πr | transcript) ≥ r log(n!)− O(c).I H(i0|π1, . . . , πr , transcript) ≥ log(n)− o(1).I H(j0|π1, . . . , πr , transcript) ≥ log(n)− o(1).I H(1[πr · · · π1(i0) = j0]|i0, π1, . . . , πr , transcript) ≥

1− o(1).I Few others...

1Technically under 12 (PVYES +PVNO).

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Conclusion

First round/communication tradeoffs for CRG &SKG for r > 2 rounds.Explicitly constructed source µ = µr ,n,L s.t.:

I ∃ efficient (i.e. O(r log n) communication) (r + 1)-roundprotocol for L-CRG/SKG

I Any (r/2− 2)-round protocol for L-CRG/SKG hascommunication Ω (minL, n).

Open questions:I Improve (r + 1)-vs-(r/2− 2) tradeoff to (r + 1)-vs-r forµr ,n,L?

I Extend analysis to amortized case?

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I am grateful to the NSF for a SODA travel grant.

Thank you!

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Proof Outline of Reduction 2

Overall goal: µ & µX × µY are indistinguishable to(r ′, c)-protocols.

Intermediate distr. µmid:

Aj = Bj forrandom j .

I.e., j isindependent ofπr · · · π1(i0).

π1

i0

π2

π3

πr-1

πr

Aj

A1

A3...A

2B

jB

3...B

2B

1


2 , π

4 , …, B

1 , …, B

n )


1 , π

3 , …, A

1 , …, A

n )

An

... Bn

...

19

Proof Outline of Reduction 2

Intermediate distr. µmid:

Aj = Bj for random j .

I.e., j isindependent ofπr · · · π1(i0).

π1

i0

π2

π3

πr-1

πr

Aj

A1

A3...A

2B

jB

3...B

2B

1


2 , π

4 , …, B

1 , …, B

n )


1 , π

3 , …, A

1 , …, A

n )

An

... Bn

...

µX × µY indist. from µmid to protocols w/communication o(n) (set disjointness hard).

PVYES & PVNO indist. to (r ′, c)-protocols⇒ µ indist. from µmid to (r ′, c)-protocols.

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Proof of indisting. of PVYES & PVNO

Idea: “round elimination” ([Nisan & Wigderson, ’93]).

“Problem” 1: Want to work with a functionalproblem.

Solution 1:Π disting. PVYES & PVNO

⇔ Π outputs 1[j0 = πr · · · π1(i0)] whp

under PVMIX = 12(PVYES + PVNO).

“Problem” 2: Players’ inputs under PVMIX aren’tindependent.

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