# Hardness from integrality gaps

UCSD winter school on Sum-of-Squares, January 2017

# Hardness and integrality gaps

### Max Cut Intergrality gap

Graph $$G$$ s.t. $$\max f_g \leq s$$ but $$\exists$$ $$d$$-pdist $$\mu$$ s.t. $${\tilde{\mathbb{E}}}_\mu f_G \geq c$$.

(Bigger $$d$$, larger $$c$$, smaller $$s$$ make gap stronger)

$$NP$$-hardness of approximation $$\Rightarrow^*$$ integrality gap.

Can we obtain the reverse implication?

A priori unlikely: What if sos is not the best algorithm?

Thm: If Khot’s Unique Games Conjecture (UGC) is true then can translate any integrality gap (even for degree $$2$$) to corresponding hardness of approximation.

Moreover, generalizes for every constraint satisfaction problem.

(Khot-Kindler-Mossell-O’Donnell, Mossell-O’Donnell-Oleszkiewicz, Raghavendra)

### Max Cut Integrality Gap (KKMO’04,MOO’05)

Thm: If there is a $$1-\epsilon$$ vs. $$1-\delta$$ gap for Max Cut, then can reduce unique games problem to distinguish given $$G$$ between (i) $$\max f_G \leq 1-\epsilon+o(1)$$ and (ii) $$\max f_G \geq 1-\delta-o(1)$$.

Proof idea: (Generalizes to any CSP (Raghavendra ’08)) ### Proof outline:

$$2LIN(k)$$ Problem: Given set $${\mathcal{E}}$$ of equations of form $$x_i + x_j = a_{i,j} (\mod k)$$, compute $${\mathop{\mathrm{val}}}({\mathcal{E}}) = \max_{x\in({\mathbb{Z}}_k)^n} \color{red}{\text{ frac. of eqs satisfied by x }}$$

$$UG(\epsilon)$$ Problem: Distinguish between $${\mathop{\mathrm{val}}}({\mathcal{E}}) \geq 1-\epsilon$$ and $${\mathop{\mathrm{val}}}({\mathcal{E}}) \leq \exp(1/\epsilon)$$ where $${\mathcal{E}}$$ is $$2LIN(k=\exp(1/\epsilon^2))$$ instance.

### Proof outline:

KKMO: Reduce $$UG(\epsilon)$$ to $$0.878 + {\mathop{\mathrm{poly}}}(\epsilon)$$ Max Cut approx.

❤ of reduction: An encoding $$E:{\mathbb{Z}}_k\rightarrow{\{0,1\}}^{2^k}$$ and graph $$\tilde{G}$$ on $$2^k$$ vertices satisfying:

Completeness: $$\forall x\in{\mathbb{Z}}_k$$, $$E(x)$$ is cut of value $$\alpha$$ in $$\tilde{G}$$.

Soundness: For every $$y\in{\{0,1\}}^{2^k}$$, if $$y$$ cuts $$\gt0.878\alpha+\epsilon$$ edges, $$y$$ is “correlated” with $$E(x)$$.

### Gap $$\Rightarrow$$ Isoperimetry $$\Rightarrow$$ Gadget

$$G=([n],E)$$ s.t.

$$\max f_G \leq 1 -\epsilon$$ but $${\tilde{\mathbb{E}}}_\mu f_G \geq 1-\delta$$

$$\Rightarrow$$

$$(X,Y)$$ over $${\mathbb{R}}^{2d}$$ s.t. $${\mathbb{E}}X_iY_i \geq 1-\delta$$ but $${\mathbb{E}}[\Psi(X)\Psi(Y)] \leq 1-\epsilon$$ $$\forall \Psi:{\mathbb{R}}^d\rightarrow {\{0,1\}}$$.

Graph $$\tilde{G}=({\{0,1\}}^d,\tilde{E})$$ s.t. $$\max f_{\tilde{G}} \geq 1-\delta$$ max cuts are $$\Psi_i(x)=x_i$$ ($$i=1..n$$).

$$\Leftarrow$$

$$(X,Y)$$ over $${\{0,1\}}^{2d}$$ s.t.

$${\mathbb{E}}X_iY_i \geq 1-\delta$$

Soundness of gadget: If $${\mathbb{E}}_{\{x,y\}\in \tilde{E}} (\Psi(x)-\Psi(y))^2 \geq 1 - \epsilon + \Omega(1)$$ then $$\Psi \approx \Psi_i$$ in sense of having influential coordinate.

### Invariance Principle (MOO’05)

Central Limit Theorem: If $$\Psi:{\mathbb{R}}^d\rightarrow{\mathbb{R}}$$ is a linear function with $$\Psi_i^2 \ll {\| \Psi \|}^2$$ for all $$i$$ then

$\Psi(X_1,\ldots,X_d) \approx \Psi(B_1,\ldots,B_d)$
where the $$X_i$$’s are i.i.d Gaussians in $$N(1/2,1/4)$$ and the $$B_i$$ are i.i.d Bernoullis in $${\{0,1\}}$$.

Invariance Principle: If $$\Psi:{\mathbb{R}}^d\rightarrow{\mathbb{R}}$$ is a constant degree polynomial with $$\Psi_i^2 \ll {\| \Psi \|}^2$$ for all $$i$$ then

$\Psi(X_1,\ldots,X_d) \approx \Psi(B_1,\ldots,B_d)$
where the $$X_i$$’s are i.i.d Gaussians in $$N(1/2,1/4)$$ and the $$B_i$$ are i.i.d Bernoullis in $${\{0,1\}}$$.

$$\Psi_i^2 := {\mathbb{E}}_{X_1\ddots X_{i-1}X_{i+1}\ddots X_d} Var_{X_i} \Psi(X_1,\ldots,X_d)$$

Can drop degree assumption by “noising up” $$\Psi$$.

### Invariance Principle (MOO’05)

Invariance Principle: If $$\Psi:{\mathbb{R}}^d\rightarrow{\mathbb{R}}$$ is a constant degree polynomial with $$\Psi_i^2 \ll {\| \Psi \|}^2$$ for all $$i$$ then

$\Psi(X_1,\ldots,X_d) \approx \Psi(B_1,\ldots,B_d)$
where the $$X_i$$’s are i.i.d Gaussians in $$N(1/2,1/4)$$ and the $$B_i$$ are i.i.d Bernoullis in $${\{0,1\}}$$.

Corollary: Gadget $$\tilde{G}$$ is sound.

### Reduction outline

Transform UG instance $$\mathcal{I}$$ on $${\mathbb{Z}}_k^n$$ into max-cut instance $$\mathcal{G}$$ on $$n$$ blocks of $$2^k$$ vertices. . . . ($$\mathcal{G}$$ has many copies of $$\tilde{G}$$ on pairs of blocks.)

Completeness: $$\Rightarrow$$ If $$x\in\Sigma^n$$ is “good” then $$E(x_1)\cdots E(x_n) \in {\{0,1\}}^{n2^d}$$ “good” cut for most copies.

Soundness: $$\Rightarrow$$ If $$\Psi:{\{0,1\}}^{n2^d}\rightarrow{\{0,1\}}$$ is often “good” cut then can decode $$\Psi$$ to “good” assignment $$x$$.

### Conjectured complexity: 3SAT (modulo ETH)

### Conjectured complexity: 3XOR (modulo ETH)

### Conjectured complexity: Max Cut (modulo ETH, UGC, and subexp algorithm)

### Summary

Assuming the UGC (basic variants of) the SOS program is optimal for a great many approximation problems.

Finite obstruction for single algorithm $$\Rightarrow$$ universal hardness result.

Is the UGC true? Big open question.

### Drawbacks of results we saw:

• Integrality gaps only for degree two sos.
• Hardness results only based on unique games conjecture.
• Huge quantitative losses :blowup of $$\exp(k)=\exp(\exp({\mathop{\mathrm{poly}}}(1/\epsilon)))$$ from $$UG(\epsilon)$$ to get $$0.878+{\mathop{\mathrm{poly}}}(\epsilon)$$ hardness for max cut).

$$\epsilon \sim 0.01$$ for non trivial results, reduction can’t output graphs smaller than $$\approx 2^{2^{10^4}}$$ vertices.

### Stronger results for other problems:

• Linear degree CSP integrality gaps for sos (Grigoriev’01,Schoenebeck’08,Tulsiani’09,B-Chan-Kothari’14,…)

• NP hardness results for certain CSP’s (Håstad’95,…,Chan’13)

Interestingly, these almost match up.

Bigger open question: Tight (in approx value and running time) algorithms and NP hardness results for all CSP’s.

Conjecture: SOS Algorithm is (up to $$\pm \epsilon$$) optimal approx algorithm for all CSP’s.

UGC implies partial variant of this conjecture, could be true independently of UGC.

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