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\begin{document}
\date{} \flushbottom
\title{The Query Complexity of Program Checking by Constant-Depth
Circuits}
\author{ V. Arvind\thanks{Institute of Mathematical Sciences, Chennai
600113, India (email: arvind@imsc.ernet.in)}\\ \and
K.~V.~Subrahmanyam \thanks{Chennai Mathematical Institute,
Chennai 600 017, India (email: kv@cmi.ac.in)}\\
\and N.~V.~Vinodchandran \thanks{University of Nebraska-Lincoln,
Lincoln NE 68588-0115, U.S.A (email: vinod@cse.unl.edu). Work done
while at Institute of Mathematical Sciences, Chennai and BRICS,
Aarhus.} }
\maketitle
\begin{abstract}
We consider deterministic and randomized \acz\ program checkers for
standard P-complete and $\nco$-complete problems. Our focus is on
the {\em query complexity} of the checker: i.e.\ the number of
queries made by the checker to the program. We show that
$\Omega(n^{1-\epsilon})$ is a lower bound on the query complexity of
deterministic \acz\ checkers for these problems, for each
$\epsilon>0$, and inputs of length $n$. On the other hand, we design
randomized \acz\ checkers of {\em constant} query complexity for
these problems.
\end{abstract}
\section{Introduction}\label{prelim}
In this paper, we study program result checking (in the sense of Blum
and Kannan~\cite{BK}) with $\acz$ circuits as checkers and we focus on
the number of queries made by the checker to the checked program. We
term this parameter as the {\em query complexity} of the checker for
the given problem. The query complexity appears to be an important
parameter in the design of efficient program checkers, because a large
query complexity can be a serious bottleneck for a checker that may
otherwise be efficient.
We first formally define program checkers introduced in~\cite{BK}.
\begin{definition}{\rm\cite{BK}}
Let $A$ be a decision problem, a program checker for $A$, $C_A$, is
a (probabilistic) oracle algorithm that, given as oracle any program
$P$ (supposedly for $A$) that halts on all instances, and for any
input instance $x$ of $A$, and any positive integer $k$ (the
security parameter) presented in unary the following holds:
\begin{enumerate}
\item
If $P$ is a correct program, that is, if $P(x)=A(x)$ for all instances
$x$, then with probability $1$, $C^P_A(x,k)$=Correct.
\item
If $P(x)\not=A(x)$ then with probability $\geq 1-2^{-k}$,
$C^P_A(x,k)$=Incorrect.
\end{enumerate}
The probability is computed over the sequences of coin flips that
$C_A$ could have tossed.
\end{definition}
When we speak of \acz\ checkers we mean that the checker $C_A$ is
described by a (uniform) family of \acz\ circuits $\{C_n\}_{n\geq 0}$,
where $C_n$ is the checker for input size $n$. We will also consider
the (stronger) notion of deterministic checkability. The decision
problem $A$ is said to be {\em deterministically} checkable if $C_A$
in the above definition is a deterministic algorithm. Next we define
the query complexity of \acz\ checkers.
\begin{definition}
Let $L$ be a decision problem that is \acz\ checkable. The \acz\
checker defined by the circuit family $\{C_n\}_{n\geq 0}$ is said to
have {\em query complexity} $q(n)$ if $q(n)$ bounds the number of
queries made the checker circuit $C_n$ for any input $x\in\Sigma^n$.
\end{definition}
The seminal paper of Blum and Kannan~\cite{BK} already initiates the
study of parallel checkers. They give a deterministic CRCW PRAM
constant-time (i.e.~\acz) program checker for the P-complete problem
{\it LFMIS} (lex.~first maximal independent set problem for graphs).
Rubinfeld in ~\cite{Rub} makes a comprehensive algorithmic study of
parallel program checkers: parallel checkers are designed
in~\cite{Rub} for various problems with emphasis on analyzing parallel
time and processor efficiency. In particular, an \acz\ checker for the
P-complete problem of evaluating straight-line programs is described
in~\cite{Rub}. However, in the context of the present paper, we note
that the above-mentioned \acz\ checkers for P-complete problems
described in~\cite{BK,Rub} have large query complexity (the number of
queries is proportional to the input size). Indeed the \acz\ checkers
in~\cite{BK,Rub} are deterministic, and as shown in the present paper,
deterministic \acz\ checkers for standard P-complete problems must
necessarily have large query complexity.
Regarding query complexity, we note that constant query checkers are
mentioned in~\cite{AHK} as a notion of practical significance. For
instance, it is shown in~\cite{AHK} that GCD has a constant query
checker.
Blum, Luby, and Rubinfeld, in their paper on self-testing and
self-correction \cite{BLR}, define and study following notions of
efficiency of self-testers and correctors. These notions are also
applicable to program checkers. We state their definition as applied
to program checkers and discuss how it is related to the query
complexity of checkers.
\begin{definition}{\rm\cite{BLR}}
Let $A$ be a decision problem and $C_A$ be a program checker for
$A$. Given a purported program $P$ for $A$, and input $x$ the
\emph{incremental time} is the running time of $C_A$ not counting
the time for oracle calls to the program $P$. The \emph{total time}
of the checker $C_A$ given input $x$ and program $P$ is the running
time of $C_A$ including the time for calls to the program $P$. The
checker $C_A$ is called \emph{different} if the incremental time is
faster than the running time for any correct program for $A$. The
checker $C_A$ is called \emph{efficient} if the total time is linear
in the running time of $P$.
\end{definition}
In the above definition of efficiency the number of queries made by
$C_A$ to $P$ is not directly measured. It is interesting to note that
even when the checker $C_A$ is both \emph{incremental} and
\emph{efficient}, it could be making a non-constant number of queries
to $P$.
Apart from these related studies, to the best of our knowledge, there
is no explicit mention of query complexity of program checkers in the
existing literature.
Our motivation for studying $\acz$ checkers is two-fold. First, in a
complexity-theoretic sense \acz\ represents the easiest model of
parallel computation, and one aspect of program checking is to try and
make the checker as efficient as possible. In this sense \acz\
checkers can be seen as constant-time parallel checkers since they
essentially correspond to constant-time CRCW PRAM algorithms. The
second motivation rests on the main goal of this paper, namely, to
study the query complexity of program checkers. It turns out that
\acz\ yields a model of program checking that is amenable to lower
bound techniques. The problems we consider are different suitably
encoded versions of the {\em Circuit Value Problem} (henceforth CVP).
Different versions of the CVP are known to be complete for different
important complexity classes: the unrestricted version is P-complete
and the CVP problem for circuits that are formulas (i.e. fanout of
each gate is one) is $\nco$-complete.
We obtain nontrivial lower bounds on the query complexity of
deterministic \acz\ checkers for these circuit value problems: we show
the lower bounds by first noting for the \Parity\ problem that, by a
result of Ajtai~\cite{Aj}, $\Omega(n^{1-\epsilon})$ is a lower bound
on the query complexity of deterministic \acz\ checkers for \Parity.
Since Parity is \acz\ many-one reducible to each considered
circuit-value problem, the same lower bounds on query complexity can be
shown to carry over. Thus, we get that for each $\epsilon>0$,
$\Omega(n^{1-\epsilon})$ is a lower bound on the query complexity of
deterministic \acz\ checkers for these circuit-value problems.
In contrast, we design randomized \acz\ checkers of {\em constant}
query complexity for these circuit-value problems. We outline the
ideas involved: consider a deterministic \acz\ checker for the given
circuit-value problem (e.g. the one described in~\cite{Rub} for
general straight-line programs). The query complexity of this checker
is roughly the number of gates in the input circuit. Our randomized
{\em constant query} \acz\ checkers for the circuit-value problems use
ideas from the PCP$(n^3,1)$ protocol for satisfiability~\cite{almss}.
Intuitively, it turns out that the number of probes into an NP proof
by a PCP protocol corresponds to the number of queries that the
checker needs to make for a given instance of the circuit-value
problem. A difficulty in the checker setting (which does not arise in
the PCP$(n^3,1)$ protocol) is that the \acz\ checker needs to compute
unbounded $\GF(2)$ sums. It cannot directly do this because Parity is
not computable in \acz. But we can get around this difficulty by using
Rubinfeld's parallel checker for Parity~\cite{Rub} which we note is
already an \acz\ constant-query checker. Since Parity is \acz\
many-one reducible to the considered circuit-value problems, we can
use the tested program for CVP to compute Parity and use Rubinfeld's
checker as subroutine to check that the returned answer is correct.
Thus we are able to design a constant-query \acz\ checker for CVP.
Since CVP is logspace many-one equivalent to checking feasibility of
linear programs FLP, our result also implies the existence of a
randomized $\nct$ constant-query checker for FLP. In fact, the same is
true for all P-complete problems that are logspace many-one equivalent
to CVP.
As explained in~\cite{almss}, the PCP theorem has evolved from
interactive proofs~\cite{GMR} and program checking~\cite{BK,BLR}. In
particular, ideas from the area of self-correcting programs play a
role in the PCP$(n^3,1)$ protocol for 3-SAT~\cite{almss}. It is not
surprising, therefore, that ingredients of the PCP$(n^3,1)$ protocol
find application in our result on program checking. Our emphasis on
the query complexity of program checkers leads naturally to ideas
underlying probabilistically checkable proofs.
%\section{Preliminaries}\label{prelim}
We now describe the CVP problems and the encodings of their instances.
Let $C$ denote a boolean circuit over the standard base (of NOT, AND,
and OR gates). We consider circuits of fanin bounded by two. We will
encode the circuit $C$ as 4-tuples $(g_1,g_2,g_3,t)$ where $t$ is a
constant number of bits to indicate the type of the gate labeled
$g_1$, and $g_2$ and $g_3$ are the gates whose values feed into the
gate labeled $g_1$. For uniformity, we can assume that NOT gates are
also encoded as such 4-tuples, except that $g_2=g_3$. Furthermore, we
insist that in the encoding, the gate labels $g_i$ be {\em
topologically sorted} consistent with the {\it DAG} underlying the
circuit $C$. Thus, in each 4-tuple $(g_1,g_2,g_3,t)$ present in the
encoding of $C$, it will hold that $g_1>g_2$ and $g_1>g_3$. This
stipulation ensures that checking whether an encoding indeed
represents a circuit can be done in \acz. For a circuit $C$ with $n$
inputs let $C_g(x_1,x_2,\dots,x_n)$ denote the value of the circuit
$C$ at gate $g$ designated as the output gate. We now define the {\em
circuit value problem} which is the decision problem that we shall
be mainly concerned with in this paper: $\{(C,g,x_1,x_2,\dots,x_n)\mid
C_g(x_1,x_2,\dots,x_n)=1\}$. The circuit $C$ is encoded as described
above.
The above circuit value problem is known to be P-complete under
projection reducibility (see \cite{GHR} for details). We denote this
by \CVP. If the input circuit is a formula (i.e.~each gate of the
input circuit has fanout at most $1$) the corresponding circuit value
problem is known to be $\nco$-complete under projection reducibility
(see \cite{GHR}); we denote this problem by $\FVP$ (for formula value
problem). Notice that an \acz\ circuit can check if a given circuit is
a formula or not. Finally, we make another important stipulation on
the circuits that are valid inputs for all the circuit value problems
that we consider: we insist that the {\em fanout} of each gate is
bounded by two. Notice that this last restriction on the input
circuits does not affect the fact that \CVP\ remains P-complete (in
fact, such a restriction already holds for the $\CVP$ in the standard
P-completeness proof by simulating polynomial-time Turing machines).
Also, observe that this extra stipulation on the input circuits can be
easily tested in \acz.
\section{Deterministic $\acz$ checkers}\label{search}
We first recall deterministic \acz\ checkers for Parity and the
circuit value problem. These are essentially implicit
in~\cite{BK,Rub}, although not in the form stated below in which we
emphasize the query complexity $n/\log^{O(1)}n$.
\begin{proposition}\label{upper1}
For each constant $k$, $\Parity(x_1,x_2,\dots,x_n)$ has a
deterministic $\acz$ checker of query complexity $n/\log^{k}n$.
\end{proposition}
\begin{proof}
Let $P$ be an alleged program for the Parity function. First
observe that the \acz\ checker can make parallel queries to $P$ for
$\Parity(x_1,x_2,\dots,x_i)$ for $2\leq i\leq n$. In order to verify
that the program's value of $\Parity(x_1,x_2,\dots,x_n)$ is correct
the checker just has to verify that the answers to the queries for
$\Parity(x_1,x_2,\dots,x_i)$ are all locally consistent:
$P(x_1,x_2,\dots,x_{i+1})=x_{i+1}\oplus P(x_1,x_2,\dots,x_i)$ for
$2\leq i\leq n-1$. The above verification can be easily done in
parallel in \acz\ since query answers $P(x_1,x_2,\dots,x_i)$ for
$2\leq i\leq n$ are available. This yields an \acz\ checker which
makes $n-1$ queries.
In order to design a checker with the number of queries scaled down
to $n/\log^{k}n$ notice that we can compute the parity of $\log n$
boolean variables in \acz\ by brute force. Thus, we can group the
$n$ input variables into $n/\log n$ groups of $\log n$ variables
each, compute the parity of each group again by brute force, and the
problem boils down to checking the program's correctness for the
parity of $n/\log n$ variables, which we can do as before with
$n/\log n$ queries to the program. To reduce the query complexity to
$n/\log^{2}n$, we iterate the above process. We group the $n/\log
n$ variables whose parity is to be checked, into groups of $\log n$
variables and compute the parity of each such group of $\log n$
variables in \acz. We are then left with the task of checking the
program's correctness for the parity of $n/\log^2 n$ variables,
which we can do as before with $n/\log^2 n$ queries to the program.
Iterating this process $k$ times we get a query complexity of
$n/\log^{k}n$. However in the process of reducing query complexity,
the depth and size of the \acz\ checker increases.
\end{proof}
\noindent
{\bf Remark}:~Notice that the above result applies to checking
iterated products over arbitrary finite monoids. The proof and
construction of the checker is similar.
Indeed, the above idea of doing piece-wise ``local testing'' in order to
check global correctness is already used in the deterministic CRCW
PRAM checker in~\cite{BK} for an encoding of the P-complete problem
{\it LFMIS}, and also in~\cite{Rub} for checking straight-line
programs. Notice that the checker for general straight-line programs
includes checking the circuit-value problem as a special case. We
include a proof sketch below since it is the starting point for the
main results of this paper.
\begin{theorem}{\rm\cite{Rub}}\label{upper2}
\CVP\ has a deterministic \acz\ checker.
\end{theorem}
\begin{proof}
Let $P$ be an alleged program for \CVP\ and let
$(C,g,x_1,\dots,x_n)$ be an input instance for program $P$. The
\acz\ checker first queries in parallel the program for
$P(C,g,x_1,\dots,x_n)$ for all gates $g\in C$. Next, for each tuple
$(g_1,g_2,g_3,t)$ in the circuit description $C$ the checker
verifies that the program's answers are consistent with the gate
type. This is again done in parallel for each tuple. The checker
must also validate the input by verifying that the tuples that
describe the circuit indeed describe an acyclic digraph. This is
made sure as described in our encoding of the instances of the \CVP.
It suffices to check that $g_1>g_2$ and $g_1>g_3$ for each tuple
$(g_1,g_2,g_3,t)$, which can be done in \acz. This completes the
proof.
\end{proof}
It can be shown similarly that $\FVP$ has a deterministic \acz\
checker. We now turn to lower bounds on the query complexity of \acz\
checkers for \CVP\ and $\FVP$. We first observe the following
property of languages $L$ having deterministic \acz\ checkers.
\begin{lemma}\label{nondet}
Let $L$ be a decision problem that is deterministically \acz\
checkable and has an \acz\ checker of query complexity $q(n)$. Then,
for each $n>0$, there is a nondeterministic \acz\ circuit that takes
$n$ input bits and $q(n)$ nondeterministic bits and accepts an input
$x\in\Sigma^n$ iff $x\in L^{=n}$ (where $L^{=n}$ denotes the set of
strings in $L$ of length $n$.)
\end{lemma}
The proof of the above lemma is clear from the definition of checking and
so we omit it.
Observe that, by symmetry, such nondeterministic \acz\ circuits also
exist for $\overline{L}$. We next recall a result due to Ajtai~\cite{Aj} on
lower bounds for \acz\ circuits approximating Parity.
\begin{theorem}{\rm\cite{Aj}}\label{ajtai}
For all constants $k, c$, and $\epsilon>0$, there is no depth $k$
circuit of size $n^c$ that can compute $\Parity(x_1,x_2,\dots,x_n)$
for more than a $1/2+2^{-n^{1-\epsilon}}$ fraction of the inputs.
Thus, no nondeterministic \acz\ circuit with $O(n^{1-\epsilon})$
nondeterministic bits can compute $\Parity(x_1,x_2,\dots,x_n)$.
\end{theorem}
\begin{lemma}\label{det-parity}
\Parity\ does not have deterministic \acz\ checkers that make
$O(n^{1-\epsilon})$ queries for any $\epsilon>0$.
\end{lemma}
\begin{proof}
Assume that \Parity\ has \acz\ checkers that makes
$O(n^{1-\epsilon})$ queries for some $\epsilon>0$. From
Lemma~\ref{nondet} it follows that for each $n>0$, there is a
nondeterministic \acz\ circuit that takes $n$ input bits and
$O(n^{1-\epsilon})$ nondeterministic bits and accepts $x\in\Sigma^n$
iff $x$ has odd parity. This is impossible since it contradicts
Ajtai's result (Theorem~\ref{ajtai} above).
\end{proof}
\begin{theorem}\label{det-cvp}
\CVP\ (likewise $\FVP$) does not have deterministic \acz\ checkers
that make $O(n^{1-\epsilon})$ queries for any $\epsilon>0$.
\end{theorem}
\begin{proof}
We prove it just for \CVP. Notice that we can design an \acz\
circuit (call it $C'$) such that given an instance $x\in\Sigma^n$ of
\Parity\ the \acz\ circuit produces an instance
$(C,x_1,x_2,\dots,x_n,g)$ of \CVP\ such that
$\Parity(x_1,x_2,\dots,x_n)=1$ iff it holds that
$(C,x_1,x_2,\dots,x_n,g)\in\CVP$. Moreover, the size of
$(C,x_1,x_2,\dots,x_n,g)$ is $O(n\log n)$, since $C$ encodes the
linear-sized circuit for \Parity\ in the 4-tuple encoding we are
using for \CVP\ instances. Assume that \CVP\ has an \acz\ checker
that makes $O(n^{1-\epsilon})$ queries for some $\epsilon>0$.
Combining the nondeterministic circuit given by Lemma~\ref{nondet}
with the \acz\ circuit $C'$, we get a nondeterministic \acz\ circuit
that takes $n$ input bits and $O(n^{1-\delta})$ nondeterministic
bits and accepts an input $x\in\Sigma^n$ iff $x$ has odd parity, for
some suitable $\delta>0$. This contradicts Lemma~\ref{det-parity}
and hence completes the proof.
\end{proof}
\section{A constant query randomized $\acz$ checker for CVP}
We first recall the relevant definition and results from~\cite{BLR}
concerning the linearity test.
\begin{definition}{\rm\cite{BLR}}
Let $F$ be $\GF(2)$ and $f, g$ be functions from $F^n$ to $F$. The
relative distance $\Delta(f,g)$ between $f$ and $g$ is the fraction of
points in $F^n$ on which they disagree. If $\Delta(f,g)\leq \delta$ then
$f$ is said to be $\delta$-close to $g$.
\end{definition}
\begin{theorem}{\rm\cite{BLR}}\label{lin-test}
Let $F$ be $\GF(2)$ and $f$ be a function from $F^n$ to $F$ such
that when we pick $y,z$ randomly from $F^n$,
$Prob[f(y)+f(z)=f(y+z)]\geq 1-\delta$, where $\delta<1/6$. Then $f$
is $3\delta$-close to some linear function.
\end{theorem}
The theorem gives a linearity test that needs to evaluate $f$ at only
a constant number of points in $F^n$, where the constant depends on
$\delta$. If $f$ passes the test then the function is guaranteed to be
$3\delta$-close to some linear function. Given a function $f$
guaranteed to be $\delta$-close to a linear function $\hat{f}$, and
$x$ in the domain of $f$, let us denote by $\mbox{SC-}f(x)$ the value
$f(x + r) - f(r)$, for a randomly chosen $r$ in the domain of $f$.
Then the above theorem guarantees that with high probability
$\mbox{SC-}f(x)$ is equal to $\hat{f}(x)$.
Next we recall another lemma from \cite{almss} (as stated in
\cite[Lemma 7.13]{MRbook}). Given $\vec{a}$ in $\GF(2)^n$, the {\em outer
product} $\vec{b}=\vec{a}\otimes\vec{a}$ is an $n\times n$ matrix
over $\GF(2)$ such that $\vec{b}_{ij}=\vec{a}_i\vec{a}_j$.
\begin{lemma}{\rm{\cite{almss}}}\label{tensor}
Let $\vec{a}\in\GF(2)^n$ and $\vec{b}$ be an $n\times n$ matrix over
$\GF(2)$. Suppose $\vec{b}\neq\vec{a}\otimes\vec{a}$, then
$\prob[\vec{r}^t(\vec{a}\otimes\vec{a})\vec{s}\neq
\vec{r}^t\vec{b}\vec{s}]\geq 1/4$
where $\vec{r}$ and $\vec{s}$ are randomly chosen from
$\GF(2)^n$.
\end{lemma}
A randomized $\acz$ checker for Parity is sketched in~\cite{Rub} based
on the fact that Parity is random self-reducible. Additionally, we
observe that this checker has constant query complexity and we will
use it as subroutine in the checker for CVP. We can also design a
constant query $\acz$ checker for Parity (without using random
self-reducibility). This checker is very similar to part of the Graph
Isomorphism checker. Since our checker is entirely different (and
\cite{Rub} gives only a sketch), we include a proof here to make the
paper self-contained.
\begin{theorem}{\rm\cite{Rub}}\label{par-const}
\Parity\ has a randomized $\acz$ checker of constant query
complexity.
\end{theorem}
\begin{proof}
Our $\acz$ checker is very similar to the standard IP protocol for
Graph Nonisomorphism. Let $P$ be a purported program for $\Parity$
and $x$ be an input from $\{0,1\}^n$. Under coordinate-wise GF$(2)$
addition, notice that $\{0,1\}^n$ is an abelian group of size $2^n$.
The set $\even$ of even parity vectors in $\{0,1\}^n$ forms a
subgroup of size $2^{n-1}$. Now, observe that $\even$ has a standard
group action on the whole of $\{0,1\}^n$: $\vec{a}\in\even$ maps
$\vec{x}\in\{0,1\}^n$ to $\vec{a}+\vec{x}$. Fix two vectors
$E=0^n\in\even$ and $O=10^{n-1}\not\in\even$. The two equivalence
classes of $\{0,1\}^n$ under this action are the sets $\even$ and
$\even+O$ of even parity vectors and odd parity vectors
respectively.
Notice that the input $x$ has even parity iff $E$ and $x$ are
$\even$-equivalent, and $x$ has odd parity iff $O$ and $x$ are
$\even$-equivalent.
Now, let $P$ be a purported program for parity. If the output $P(x)$
is ``odd'', it is the same as the program asserting that $x$ and $E$
are not $\even$-equivalent. We can check this by making a single
query with probability $1/2$ of catching the error as follows: pick
either $x$ or $E$ at random and call it $\vec{y}$. Pick a random
element $\vec{a}\in\even$ and compute $\vec{a}+\vec{y}$. The checker
queries the program for $\vec{a}+\vec{y}$ and outputs ``pass'' if
and only if $P$ answers ``odd'' if $y=x$ and ``even'' if $y=E$.
Notice that this is exactly like the standard IP protocol for Graph
Nonisomorphism.
If the output $P(x)$ is ``even'' then the checker follows the same
algorithm as above with $O$ replacing $E$ and other appropriate
changes.
It remains to describe a randomized $\acz$ algorithm for generating
elements of $\even$ uniformly at random. Seen as a vector space,
$\even$ has a basis consisting of the following $n-1$ vectors, $b_1
= \{1,1,0,\ldots,0\}$,$b_2 = \{0,1,1,0,\ldots,0\}$,$\ldots$,$b_{n-1}
= \{0,0,\ldots,1,1\}$. In order to generate a random element of
$\even$, the checker picks up a random vector $k$ in $\{0,1\}^{n-1}$
and computes $\Sigma_i k_ib_i$. Although this involves a parity
computation for each coordinate, notice that for each coordinate
there are at most two non-zero terms (in known positions) in the
parity sum to be computed. Thus, such a parity computation can be
easily done in \acz.
This completes the proof.
\end{proof}
\noindent{\bf Remark.} The above result can be generalized to the
equivalence problem under group actions, where the group can be
sampled in $\acz$ and there are a constant number of equivalence
classes.\\
We are now ready to design the constant query checker for the \CVP\
problem (also for $\FVP$). We make use of ideas in the PCP$(n^3,1)$
protocol for 3-SAT from~\cite{almss}. A crucial point of departure
from~\cite{almss} is when the checker needs to compute the parity of a
multiset of input variables and a product of input variables. To do
this we use as subroutine the checker of Theorem~\ref{par-const}.
Another point to note is that all queries have to be valid instances
of \CVP\ (or $\FVP$ as the case may be), and they need to be generated
in \acz. The starting point is the deterministic \acz\ checker for
\CVP\ described in Theorem~\ref{upper2}. Recall that given instance
$(C,g,x_1,\dots,x_n)$ the deterministic \acz\ checker queries the
program for $(C,g_i,x_1,\dots,x_n)$, for each gate $g_i$ of $C$. Then
it checks that the query answers are locally consistent for each gate.
Let $y_1,y_2,\dots,y_m$ be the query answers by $P$ for the queries
$(C,g_i,x_1,\dots,x_n)$, $1\leq i\leq m$, where $C$ has $m$ gates. The
{\em unique correct} vector $y_1,y_2,\dots,y_m$ is a satisfying
assignment to the collection of all the gate conditions (each of which
is essentially a 3-literal formula). The idea is to avoid querying
explicitly for $y_i$'s. Instead, using randomness the checker will
make fewer queries for other inputs that encode the $y_i$'s. More
precisely, we need to encode $y_1,y_2,\dots,y_m$ in a way that making
a constant number of queries to the program (which is similar to a
constant number of probes into a proof by a PCP protocol) can convince
the \acz\ checker with high probability that $y_1,y_2,\dots,y_m$ is
consistent with all the gate conditions.
\begin{theorem}\label{rand-cvp}
The $\p$-complete problem \CVP\ (likewise the $\nco$-complete
problem $\FVP$) has a randomized \acz\ checker of constant query
complexity.
\end{theorem}
\begin{proof}
We describe the checker only for \CVP\ (the checker for $\FVP$ is
similar). Let $P$ be a program for \CVP\ and $(C,g,x_1,\dots,x_n)$
be an input. Suppose C has $m$ gates $g_1,\dots,g_m$ w.l.o.g.~assume
$g_m=g$. The deterministic checker of Theorem~\ref{upper2} queries
$P$ for $(C,g_i,x_1,\dots,x_n)$, for each $g_i$. It then performs a
local consistency check to test the validity of
$(C,g_m,x_1,\dots,x_n)$. The new checker must avoid querying
explicitly for each $y_i := (C,g_i,x_1,\dots,x_n)$. Let $t_i$
denote a $\GF(2)$ polynomial corresponding to the $i$th gate of the
circuit $C$, where $t_i$ is a polynomial in at most three variables
(these three variables are from
$\{x_1,\dots,x_n\}\cup\{y_1,y_2,\dots,y_m\}$). Define $t_i$ such
that it is zero iff the corresponding variables are consistent with
the gate type of $g_i$. More precisely, let $g$ be an AND gate with
output $a$ and inputs $b$ and $c$ then the polynomial corresponding
to $g$ is $a+bc$. Similarly, for an OR gate with output $a$ and
inputs $b$ and $c$ the polynomial is $a+b+c+bc$, and for a NOT gate
with output $a$ and input $b$ it is $a+b+1$.
\vspace{3mm}
\noindent {\bf The checker}
The checker first computes $y_m = P(C,g,x_1,\dots,x_n)$ by querying
the program on input $(C,g,x_1,\dots,x_n)$. It suffices for the \acz\
checker to verify that the following linear function
$f(\vec{x},\vec{y},\vec{z})=\Sigma_{i=1}^m t_i(\vec{x},\vec{y})z_i$ in
the new variables $z_i, 1\leq i\leq m$ is the zero linear function.
Notice that if this function is nonzero then the linear function
$f(\vec{x},\vec{y},\vec{z})$ evaluated at a randomly chosen
$\vec{z}:=\langle z_1,\ldots,z_m\rangle$ would be nonzero with
probability $1/2$. If it were the zero function then it must be zero
with probability $1$.
Recall that the checker must compute this value by making a series of
\CVP\ queries that are generated in \acz. Towards this end, we rewrite
the above expression: $f(\vec{x},\vec{y},\vec{z})=
p(\vec{x},\vec{z})+q(\vec{y},\vec{z})+r(\vec{y},\vec{z})$, where
$p(\vec{x},\vec{z}) = \Sigma_{i=1}^{n}\Sigma_{k=1}^{m}\chi_{i}^{k} x_i
+
\Sigma_{i=1}^{n}\Sigma_{j=1}^{n}\Sigma_{k=1}^{m}\chi_{ij}^{k}x_ix_j$,
$q(\vec{y},\vec{z}) = \Sigma_{k=1}^{m}c_i*y_i$, and
$r(\vec{y},\vec{z}) = \Sigma_{(i,j)\in [m]\times[m]} c_{ij}*y_iy_j$.
Since the gates in the circuit have input at most $2$, notice that if
we write $f(\vec{x},\vec{y},\vec{z})$ as a sum of terms, then any term
is either independent of the $y_i$'s, a linear function of the
$y_i$'s, or is quadratic in the $y_i$'s. Observe that we have,
accordingly, written $f(\vec{x},\vec{y},\vec{z})$ in three parts. The
first part $p(\vec{x},\vec{z})$ is independent of $\vec{y}$. The part
$q(\vec{y},\vec{z})$ is linear in the $y_i$'s, and the part
$r(\vec{y},\vec{z})$ is quadratic.
We have similarly written $p(\vec{x},\vec{z})$ as a function that is
quadratic in the $x_i$'s. Here, $\chi_{i}^{k}$ is defined to be 1 if
$x_i$ appears in the polynomial $t_k$ and if $z_k$ is 1, and defined
to be zero otherwise. Likewise $\chi_{ij}^{k}$, is defined to be 1 if
$x_ix_j$ appears in the polynomial $t_k$ and $z_k$ is 1, and defined
to be zero otherwise. It follows that an $nm + n^2m$ length Boolean
vector representing each term in $p$ can be obtained in \acz.
Next, observe that coefficients $c_i$ and $c_{ij}$ in $q$ and $r$
depend upon gates $g_i$ and $g_j$, the constant number of gates they
feed into, the elements of $\vec{z}$ corresponding to these gates and
a constant number of input bits. So computing each of these
coefficients involves computing the parity of a {\em constant} number
of Boolean variables. This can be done directly in \acz.\footnote{It
is easier to first conceive of a constant time CRCW PRAM algorithm
for this task.}
The checker will work as follows: It first picks a random $\vec{z}$.
Then, using polynomially many random bits, the checker will construct
a constant number of queries for the given program $P$. From the
answers given by $P$ the checker will compute bits $\tilde{p}$,
$\tilde{q}$, and $\tilde{r}$. To complete the checking the checker
evaluates $\tilde{p} + \tilde{q} +\tilde{r}$ and accepts $P$ as
correct on its input if and only if this sum is zero.
Finally we will argue that if $P$ is incorrect on its input then
$\tilde{p} + \tilde{q} +\tilde{r}\neq 0$ with constant nonzero
probability.
We now describe below how the checker computes $\tilde{p}$,
$\tilde{q}$, and $\tilde{r}$.
\vspace{3mm}
\noindent {\bf Computing $\tilde{p}$} Note that $p(\tilde{x},\tilde{z})$
is the parity of $nm + n^2m$ Boolean variables. As noted above the
value of these variables can be obtained in \acz\ given $\vec{x}$ and
$\vec{z}$. The checker constructs a description of a canonical circuit
for the parity of $nm + n^2m$ variables, and queries the program on
this input. Next the \acz\ checker checks the answer of $P$ using the
checker of Theorem~\ref{par-const} as subroutine. If the answer is
wrong then with high probability the subroutine checker will reject
the program as incorrect. Thus the \acz\ checker computes a value
$\tilde{p}$ which is equal to $p(\tilde{x},\tilde{z})$ with high
(constant) probability. In the process of this computation, only a
constant number of queries are made to $P$.
Notice that, unlike in~\cite{almss}, we have to deal with both
$y_1,\dots,y_m$ and $x_1,x_2,\dots,x_n$ which occur in the polynomials
$t_i$. The crucial difference between the $x_j$'s and $y_j$'s is that
$x_1,x_2,\dots,x_n$ are {\em bound} to the input values. Thus,
computing the value of $\tilde{p}$ is a {\em parity computation} which
the checker requires to get done. As explained above, this is
done using the checker of Theorem~\ref{par-const} as subroutine.\\
\noindent {\bf Computing $\tilde{q}$ and $\tilde{r}$}
To compute $\tilde{q}$ and $\tilde{r}$ the checker goes through the
following steps.
\begin{enumerate}
\item It builds a circuit $C_1$ with new inputs $r_1,r_2,\dots,r_m$
that computes the function $\Sigma_{i=1}^m y_ir_i$. Recall that
$y_i$ is the output of the $i$th gate of the input circuit $C$ on
input $x_1,x_2,\dots,x_n$. Clearly, the encoding of $C_1$ can be
generated by an \acz\ circuit from the encoding of $C$.
\item Similar to the above step, the checker builds a circuit $C_2$
with new inputs $r_{ij}, 1\leq i,j\leq m$ that computes
$\Sigma_{i=1}^m\Sigma_{j=1}^m y_iy_jr_{ij}$. It is clear that an
encoding for $C_2$ can also be generated by an \acz\ circuit.
\item The checker verifies that the program's behavior on $C_1$ is a
function that is $\delta$-close to a linear function in the
variables $r_i$, and the program's behavior on $C_2$ is a function
that is $\delta$-close to a linear function in the variables
$r_{ij}$. This can be done as described in the previous section
using Theorem~\ref{lin-test}. If either of the tests fails, the
checker rejects the program as being incorrect.
\item As in the PCP protocol the checker performs a consistency check:
Let $\Sigma_{i=1}^m\Sigma_{j=1}^m b_{ij}r_{ij}$ be the linear
function to which $C_2$ is $\delta$-close. The checker does a
constant query test, and ensures with high probability that the
matrix $b_{ij}$ is the tensor product of $\vec{y}$ with itself.
To do this the checker employs the test given by Lemma~\ref{tensor}.
For two randomly chosen vectors $r_1, r_2$ of length $m$ it verifies
that $\mbox{SC-}C_1(r_1)*\mbox{SC-}C_1(r_2) =
\mbox{SC-}C_2(r_1\otimes r_2)$
Note that the tensor product can be computed in \acz\ and the
checker needs to make a constant number of queries to $P$.
\end{enumerate}
Having performed the linearity and consistency tests the checker
evaluates $q$ and $r$ by self-correction. Let $\vec{c}$ denote the
$m$-vector $c_1,c_2,\dots,c_m$ and let $\vec{d}$ denote the $m\times
m$-vector consisting of $c_{ij}, 1\leq i,j\leq m$. The checker sets
$\tilde{q}$ to $\mbox{SC-}C_1(\vec{c})$ and $\tilde{r}$ to
$\mbox{SC-}C_2(\vec{d})$.
\vspace{3mm}
\noindent{\bf Correctness.}
If $P$ is correct for all inputs then with probability 1 the checker
will pass $P$ as correct. Suppose that $P$ is incorrect on input
$(C,g,x_1,\dots,x_n)$ and let $P(C,g,x_1,\dots,x_n) = b$. Let $F$ be
the event that the checker fails to detect the program as incorrect.
It suffices to show that $\prob[F]$ is bounded by a constant strictly
smaller than $1$. Since the (constant query $\acz$) checker has one
sided error, using standard techniques we can amplify the success
probability to any constant.
Let $h_1$ and $h_2$ be the designated output gates of circuits $C_1$
and $C_2$. Let $\hat{C_1}(r_1,\ldots,r_n)$ denote
$P(C_1,h_1,r_1,\ldots,r_n)$, and $\hat{C_2}(r_{11},\ldots,r_{nn})$
denote $P(C_2,h_2,r_{11},\ldots,r_{nn})$. Let $\delta>0$ be a small
constant to be suitably chosen later. Denote by $T$ the event that
$\hat{C_1}(r_1,\dots,r_n)$ is $\delta$-close to a unique linear
function $\Sigma_{i=1}^m\hat{y_i}r_i$ and
$\hat{C_2}(r_{11},\dots,r_{nn})$ is $\delta$-close to the unique
linear function $\Sigma_{(i,j)\in [m] \times[m]}
\hat{y_i}\hat{y_j}*r_{ij}$. Let
$\hat{y}=\langle\hat{y_1},\dots,\hat{y_m}\rangle$. Denote by $w$ the entire
string of random bits used by the checker; in particular, $\vec{z}$ is
a substring of $w$.
Since $T$ is independent of $w$, in order to bound $\prob_w[F]$ it
suffices to bound each of $\prob_w[F|T]$ and $\prob_w[F|\neg T]$ by a
constant smaller than $1$.
\begin{itemize}
\item[Case~(i)] Suppose $T$ is false. Notice that $\prob_w[F|\neg T]$
is bounded above by the probability that the checker fails to reject
$P$ in steps 3 and 4. Now, by Theorem~\ref{lin-test} and
Lemma~\ref{tensor}, this probability can be made smaller than any
constant $\epsilon>0$ for a suitably small constant $\delta$. This
is achieved by the checker repeating the tests in steps 3 and 4 a
constant number of times.
\item[(ii)] Next, suppose $T$ is true. We must now bound
$\prob_w[F|T]$. Given that $P$ is incorrect on its input, the
function $f(\vec{x},\hat{y},\vec{z})=\Sigma_{i=1}^m
t_i(\vec{x},\hat{y})z_i$ is a nonzero linear function of the
$z_i$'s. Hence, $\prob_{w}[f(\vec{x},\hat{y},\vec{z}) =1~ | ~T] =
\frac{1}{2}$. Now,
\begin{eqnarray*}\label{eq1}
\prob_{w}[F~ | ~T] & = & \prob_{w}[F ~\& ~f(\vec{x},\hat{y},\vec{z}) =1~ | ~T] +
\prob_{w}[F ~\& ~f(\vec{x},\hat{y},\vec{z}) =0~ | ~T] \\
&\leq & \prob_{w}[F ~\& ~f(\vec{x},\hat{y},\vec{z}) =1~ | ~T] + 1/2 \\
& = & \prob_w[f(\vec{x},\hat{y},\vec{z}) =1~ | ~T]*\prob_w[F~|~f(\vec{x},\hat{y},\vec{z}) =1~\& ~T] + 1/2\\
& = & 1/2*\prob_w[F~ | ~f(\vec{x},\hat{y},\vec{z}) =1 ~\& ~T] + 1/2.
\end{eqnarray*}
Now, observe that
\[
\prob_w[F~ |~f(\vec{x},\hat{y},\vec{z}) =1 ~\& ~T]
\leq \prob_w[\tilde{p}\neq p~|~T] + \prob_w[\tilde{q}\neq q~|~T] +
\prob_w[\tilde{r}\neq r~|~T].
\]
Since $\tilde{q}$ and $\tilde{r}$ are obtained by self-correction, we
have $\prob_w[\tilde{q}\neq q~|~T]\leq 2\delta$, and
$\prob_w[\tilde{r}\neq r~|~T]\leq 2\delta$. By
Theorem~\ref{par-const}, $\prob_w[\tilde{p}\neq p~|~T]\leq 1/2$. Thus,
we get $\prob_w[F~ |~f(\vec{x},\hat{y},\vec{z}) =1 ~\& ~T]\leq
1/2+4\delta$.
Putting it all together, we finally get
\[
\prob_w[F~|~T]\leq 1/2*(1/2+4\delta)~+~1/2.
\]
Choosing $\delta<1/16$ we can bound $\prob_w[F~|~T]$ by $7/8$.
\end{itemize}
This completes the proof.
\end{proof}
\section{Concluding remarks and open problems}
In this paper we have studied the query complexity of deterministic
and randomized \acz\ program checkers for some standard P-complete and
$\nco$-complete problems and have shown a big difference in the
respective query complexities for the considered problems. Notice that
it is crucial in the proof of Theorem~\ref{rand-cvp} to be working
with a circuit-value problem whose instances have the following
closure property: if $(C,g,x_1,x_2,\dots,x_n)$ is an instance of the
problem and $C$ has $m$ gates $g_1,\dots,g_m$, then the circuit $C'$
with $m$ new boolean inputs $r_1,\dots,r_m$ that computes
$\oplus_{i=1}^m g_ir_i$ by putting a parity gate on top to evaluate
the sum and fanin two AND gates to evaluate each $g_ir_i$ is an
instance of the same circuit-value problem. This property clearly
holds for the standard P-complete problem CVP and the $\nco$-complete
problem FVP.
On the other hand, consider the complexity class NL. It is known that
the circuit value problem with instances $(C,g,x_1,x_2,\dots,x_n)$
such that $C$ is a \emph{skew} circuit (i.e.\ $C$ has all fanin two
gates, and every AND gates has at least one of the $x_i$'s as input)
is NL-complete under constant-depth reductions \cite{HV}. In the light
of the above discussion, modify this problem to now consist of
instances $(C,g,x_1,x_2,\dots,x_n)$, such that $C$ has a formula on
top whose inputs are either the $x_i$'s or skew circuits with input
$x_1,x_2,\dots,x_n$. Call this circuit-value problem $\CVP_{\rm nl}$.
Clearly, $\CVP_{\rm nl}$ is NL-hard under constant-depth reductions.
Since $\NL^{\NL}=\NL$ it follows that $\CVP_{\rm nl}$ is in NL. Thus,
it is NL-complete under constant-depth reductions. From the previous
paragraph it is clear that the proof of Theorem~\ref{rand-cvp} will go
through for $\CVP_{\rm nl}$. Therefore, $\CVP_{\rm nl}$ is an
NL-complete problem with a constant query \acz\ checker.
We next give an example of a suitable constant query checker for a
P-complete problem that is not a circuit-value problem. We consider
feasibility of linear programs (FLP), which is a P-complete problem.
First we make a simple observation (which is a version of Beigel's
trick \cite{BK}).
\begin{observation}
If $\pi_1$ and $\pi_2$ are $\acz$ many-one equivalent problems and
if $\pi_1$ has a constant-query \acz\ checker then so does $\pi_2$.
If $\pi_1$ and $\pi_2$ are $\nco$ many-one equivalent and if $\pi_1$
has a constant-query $\acz$ checker then $\pi_2$ has a
constant-query $\nco$ checker.
\end{observation}
Now, $\CVP$ is $\nco$-reducible to $\FLP$~\cite{GHR}. Since $\CVP$ is
P-complete under projections, it is clear that $\CVP$ and $\FLP$ are
many-one equivalent under $\nco$ reductions. Thus, from the above
observation we get the following result.
\begin{theorem}
The $\p$-complete problem \FLP\ has a randomized $\nco$ checker of
constant query complexity.
\end{theorem}
Along similar lines it is possible to derive checkers for other
complete problems for P, NL, and $\nco$.
Finally, our main focus in this paper has been on query complexity
rather than on the size of the circuits defining the checker itself.
It would be interesting to investigate trade-offs between the
time/space complexities of the checker and its query complexity.
\vspace{3mm}
\noindent{\bf Acknowledgments.}~~ We thank Peter Bro Miltersen, Jaikumar
Radhakrishnan, and S.~Venkatesh for interesting discussions on \acz\
lower bounds for parity. For useful remarks we are grateful to Samik
Sengupta and D.~Sivakumar. We are grateful to the referees for their
comments and suggestions that have helped improve the presentation.
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