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\lecture{6}{Feb 26, 2007}{Madhu Sudan}{David Glasser}
Today's lecture is an alternate proof that Parity is not in $AC^0$, as
shown by Smolensky.
Recall: $AC^0$ has polynomially-sized circuits of constant depth, with
unlimited fan-in ANDs and ORs. Argument is roughly that parity is
complex, whereas $AC^0$ is simple. This is a more algebraic argument
than last time.
What measure of complexity can we use to make this argument? The
degree of the polynomial which computes the function. What field is
this polynomial in? The obvious candidate is $\F_2$. Then we can
have a polynomial on $n$ variables. This doesn't work too well: AND
ends up being hard ($x_1 x_2 \cdots x_n$ has degree $n$) and parity is
simple ($x_1 + \cdots + x_n$) has degree 1. That's backwards!
Instead, we'll use \emph{any} other field, or specifically any field
of characteristic not equal to 2.
There is a fundamental trick of complexity theory: instead of writing
about 0 and 1, we write about 1 and $-1$. We can map between this
with $1-2x$ or $\frac{1-y}{2}$. Or $(-1)^b$. Now, parity is a sum of
things in the 0/1 representation, but in the other representation,
it's about a product/power, which is more complicated.
We claim that if $p \colon \F^n \to \F$ computes the parity of its
input when the inputs are 0 or 1, then a polynomial $q$ defined as
$1-2p\left(\frac{1-y_1}{2}, \dots, \frac{1-y_n}{2}\right)$ compute the
product of the variables $y$ when each $y_i$ is $\pm 1$. Exercise for
the reader: show that $p$ and $q$ have the same degree.
We can use this to prove that the degree of $p$ must be at least $n$,
since we can show (apparently) that a product polynomial must actually
contain a $y_1 \cdots y_n$ term. I'm not sure why, but we're going to
prove something stronger today I think.
OK, so we've made parity complicated. Can we make AND easy? This is
why this is the Razborov-Smolensky theorem: this is Razborov's part.
It is the method of approximations. There's some fuzziness involved.
The idea is that we have a function that is almost the same as
whatever you're actually trying to model.
Lemma: If $C$ is an $AC^0$ circuit of depth $d$ and size $s$, then for
all $\epsilon$, there exists a $S \subseteq \zo^n$ and a polynomial
$p(x_1,\dots,x_n)$ of degree $D$ such that for all $x\in S$, $p(x) =
C(x)$ and $D \leq \left( \log\frac{s}{\epsilon} \right)^d$ and
$\abs{S} \geq (1-\epsilon) 2^n$.
Smolensky idea: fix an input, and replace gates probabilistically.
Here's a lemma towards this: There exists a distribution of
polynomials $p$ of degree $k=O \left( \log\frac{s}{\epsilon} \right)$
such that for all $y_1, y_2, \dots, y_n$, $\Pr[p(Y) = OR(Y)] \geq 1
-exp(-k)$, where the probability is over the polynomials.
Exact OR is defined via De Morgan as $1-\prod (1-y_i)$, but this has
degree $n$: ewww. We'd like low-degree almost-ORs. How does this
work? We'll define $Probably-OR(y_1,\dots,y_m)$ by picking $\alpha_1$
through $\alpha_m$ at random from $\F$ and look at $\sum \alpha_i
y_i$. If the inputs are all 0, then (like OR) the output is zero.
DeMillo-Lipton-Schwartz-Zippel Lemma: If $p \colon \F^n \to \F$ is a
nonzero polynomial of degree $d$, then $\Pr_{x \in S^n} [ p(x) = 0]
\leq \frac{d}{\abs{S}}$. This is basically just a matter of counting
zeros; we can get it by induction on the number of variables.
How do we apply this? Consider the polynomial $p(\alpha_i) = \sum
\alpha_i y_i$ (so the $\alpha_i$ are the variables). So when the
$y_i$ are not all zero, the probability that the output is nonzero is
at least $1-\frac{1}{\F}$. If we use $\F_3$, this is $\frac{2}{3}$,
and it only gets better. We might want to raise this to $\abs{\F}-1$
to make the answer always 0 or 1; so it ends up having degree 2 (for
$\F_3$). So this gives us a weak probabilistic OR.
How about a stronger one? Pick $\alpha^{(1)}$, $\alpha^{(2)}$,
\dots{}, $\alpha^{(k)}$ from $F^m$ independently. We'll let
probabilistic OR be the EXACT-OR of the elements $ \left( \sum
\alpha_i^{(j)} y_i \right)^2$. So this will have degree $2k$ to
compute the OR of $m$ things, where $k$ is not a function of $m$.
Given a circuit, we can replace NOT gates by a degree-1 polynomial; we
can DeMorganize AND gates; and we can replace OR gates with this
probabilistic OR polynomial. Thus we prove a lemma: given a circuit
$C$ of size $S$ and depth $d$ and some $\epsilon >0$, there exists a
distribution of polynomials $p$ of degree at most $ \left( \log
\frac{s}{\epsilon} \right)^d$ such that for every $x \in \zo^n$,
$\Pr_p [p(x) = C(x)] \geq 1-\epsilon$. i.e., with a fixed input and
picking polynomials randomly, we do well. Each OR gates gives degree
$k$, so by the top of the circuit, it's $k^d$. For this input, we can
choose $k$ large enough to shrink the error, and it gives us the value
of $k$ around $\log \frac{s}{\epsilon}$, I suppose.
Now, how do we prove our original lemma, which is about varying
inputs? Basically by choosing the set $S$ to be the ones where the
lemma we just ``proved'' holds. So we've shown that $AC^0$ is simple,
basically; now we need to show that parity isn't.
Our new lemma: parity cannot be approximated by low-degree
polynomials. So we'll prove that if $\abs{T} \geq (1-\epsilon) 2^n$
and $T \subseteq \{-1,1\}^n$, and $p$ is a polynomial of degree $D$
such that $p$ computes the product of its inputs for all $Y\in T$,
then $D \geq \Omega \left( \left( \half-\epsilon \right)\sqrt{n}
\right)$. How do we show this?
Consider all functions mapping $T$ into $\F_3$. There are
$3^{\abs{T}}$ such functions. All of these are polynomials of degree
at most 2 in each variable. Consider $q(y_1,\dots,y_n) = \sum_d c_d
\prod_i y_i^{d_i}$. Because we only care about the behavior when $y_i
= \pm 1$, then we can assume that each $d_i$ is either 0 or 1 (since
$y_i^2 = 1$).
Consider the product $y_1 y_2\cdots y_{\frac{3n}{4}}$; it is equal to
$y_1 y_2\cdots y_{\frac{3n}{4}} y_{\frac{3n}{4}+1}^2 \cdots y_n^2 =
(\prod y_i) (\prod_{i > \frac{3n}{4}} y_i)$. But this is equal to
$p(Y) (\prod_{i > \frac{3n}{4}} y_i)$; so the degree is now
$\frac{n}{4} + D$. In fact, for the general case of a function, we
can get that $\sum d_i$ is at most $\frac{n}{2} + D$. So if this $p$
exists, then all the functions have polynomials of degree at most
$\frac{n}{2} + D$. How may ways can you write these polynomials?
$\sum_{i=0}^{\frac{n}{2} + D} {n \choose i} \leq
\sum_{i=0}^{\frac{n}{2}} {n \choose i} + D {n \choose \frac{n}{2}}
\leq 2^{n-1} + D 2^n$. There aren't enough polynomials! So, more or
less, this shows that the theorem holds. (In fact, even given gates
to calculate parity mod 3, it's hard to calculate mod 2.)
\end{document}
% LocalWords: Madhu Glasser Smolensky ANDs ORs Razborov Razborov's ewww Zippel
% LocalWords: DeMillo DeMorganize