Parity helps to compute majority
Christian Engels September 06, 2019Resources
- [https://eccc.weizmann.ac.il/report/2019/073/download](Oliveira, Santhanam, Srinivasan)
Results
- Theorem 1: Let $d>5$, Majority on $n$ bits can be computed by a depth-$d$ $AC^0[\oplus]$ circuit of size $2^{\tilde O(n^{2/3 \frac 1 {d-4}})}$.
- Theorem 2: For any $d\geq 3$ Majority requires $AC^0[\oplus]$ circuits of depth-$d$ and $2^{\tilde O(n^{\frac 1 {2d-4}})}$ size.
- The upper bound is stronger than known $AC_0$ (without parity) lower bounds, meaning parity helps compute majority.
Proof Ideas
Upper Bound
- For inputs of hamming distance $t$ of $n/2$ use a $t$-degree polynomial to compute majority. It is symmetric and over $\mathbb{F}_2$ and can be represented as a circuit of size $\exp(t^{2/d} \log n)$.
- When the weight is $t$ far from $n/2$ use sampling and Coin Problem to get size $\exp((n/t)^{1/d})$, both of depth $d$.
Lower Bound
- General idea: any such circuit can be approximated by low degree polynomials and composing these approximations together. Then show that majority does not have low degree polynomial approximation.
- They improve the best known bound for this technique in the following way:
- The standard approximation is one-sided, meaning that the approximation is much better on $C^{-1}(0)$ versus $C^{-1}(1)$.
- Hence, they need to show that any polynomial that approximates correctly an $\epsilon$ fraction of the 0-inputs but may error on a constant fraction of the 1-inputs.
- The last part uses the combinatorics on hilbert functions.
Proof of Upper Bound
- Notice that every symmetric boolean function just looks at the weight of the assignment
- Then we just need to be able to compute the function that is one on weight $i$ and zero on every other weight ($E_i$) as $f$ can be written as a disjunction of $n+1$ different $E_i$ for different $i$.
- Then we can $D_{i,j}$ that is 1 on hamming weight $i$ and 0 on hamming weight $j$ as $E_i=\wedge_{0\leq j\leq n, i\neq j} D_{i,j}$.
- Hence, computing $D_{i,j}$ with $AC_0[\oplus]$ is enough.
- Now the proof splits into parts
\lvert i - j \rvert \leq n^
- There exists integers $c_1,\dots, c_{l-1}$ and a polynomial of degree at most $l-1$ and integer coefficients such that for all $x$, $x$ has hamming weight $k+t$, $p(x)=c_t$
- This polynomial can be constructed by a sum over elementary symmetric polynomials.
- Since parity of an integer is a ring homomorphism we can use the above lemma to compute $D_{i,j}$
- Now as this can be computed by a ABP, we can divide and conquer this ABP to construct a circuit with the top gate being a parity gate.
- This gives us: $AC_0[\oplus]$ circuit of size $n^{O(\lvert i - j\rvert^{2/d'} )}$ and depth $d'$.
$\lvert i - j \rvert \geq n^{1/3}$
- Let $Prom_{a,b}$ be the function that is 0 if the hamming weight is smaller than $a$ and 1 if the hamming weight is larger than $b$.
- $D_{i,j}$ can be given by $\Prom_{a,b}$ for with some padding of the input.
- This can now be computed by a randomized $AC_0$ circuit and this construction can be derandomized.
- Resulting in a circuit of the size: $\exp(O(n/\lvert i - j \rvert )^{1/(d'-2)})$.
Linear Threshold function for combining these two
- Any treshold function can be comp0uted by a polynomial size circuit
- This works via standard chinese remainder construction
Proof of Lower Bound
- We call a probabilistic $(\epsilon_1, \epsilon_b)$-error polynomial for a boolean function $f$ if the probability that $P(a)$ differs from $f(a)$ on $a\in f^{-1}(b)$ is bounded by $e_b$ for a bit $b\in {0,1 }$.
- Here $P$ is chosen from a distribution.
- By the Razborov construction we get polynomials that approximate the boolean function
- For the output gate, we construct instead a $(0,1/20)$-error probabilistic polynomial. Again this comes from the Razborov construction.
Lower bound
- All polynomials that have error characteristic $1/10,\epsilon$ or $\epsilon,1/10$ have degree at least $\sqrt{n\log (1/\epsilon)}$.
- A non-zero certifying polynomial $R$ for a boolean function $f$, if $f$ is constant on the support of $f$.
- If $R$ is a certifying polynomial for Majority then $\deg(R) \geq \lceil n/2 \rceil$.
- There is a non-zero degree polynomial $Q$ of low degree such that vanishes on all points in $E_0 = { x\in Maj^{-1}(0)\mid P(x)\neq 0}$.
- Let $P$ be the polynomial approximating Majority.
- Look at now $R=P\cdot Q$. If this is non-zero then it is a certifying polynomial for majority.
- Indeed we can find an $a\in supp(P)$ such that $Q(a)\neq 0$.
- Theorem: Fix any $E\subseteq \mathbb{F}_2^n$$ then $\frac{\lvert a\mid Q(a)=0 \forall Q \text{of degree at most $D$ that vanish on $E$}}{2^n} \leq \frac{\lvert E\rvert}{N_D}$ where $N_D$ is the number of multilinear monomials of degree $D$.
- Intuitively this says that if $E$ is chosen not too large then constraining a polynomial of degree $D$ to be zero on $E$ does not restrict it too much.
Questions
- Dealing with constant depth circuit, why not have the other gates be similarly "skew". With the constant depth, the error shouldn't be too large?
- This would not give a better bound as the degree lower bound already is independent of where the one-sidedness occurs.