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\lecture{11 --- March 10, 2005}{Spring 2005}{Prof.\ Erik Demaine}{Yoyo Zhou}
\section{Overview}
In this lecture, we discuss lower bounds on the cell-probe complexity
of the static predecessor problem. The parameters of interest are $w$,
the size of the machine word, and also the size of input integers, and
$n$, the set size. Observe that a requirement on space is essential
for static lower bound, since with space $2^w$ one can precompute all
answers, and respond in constant time.
In this lecture, we concentrat on bounds that are purely in $w$ or
$n$, and we assume the space is $n^{O(1)}$. In particular, we use a
technique called round elimination to show that for all $w$ there
exists an $n$ such that the problem requires $\Omega(\frac{\lg w}{\lg
\lg w})$ time, and for all $n$ there exists a $w$ requiring
$\Omega(\sqrt{\frac{\lg n}{\lg \lg n}})$ time. In the next lecture, we
will discuss the tradeoff between $w$, $n$, and space.
\section{Lower Bound Results}
We only consider the cell-probe model, and the static problem. We
require space to be polynomial in $n$. Since a static data structure
can be constructed through $n$ insertions, this implies the same lower
bound on query for the dynamic problem, which holds even if updates
take $n^{O(1)}$ time.
The first superconstant bound, proved by Ajtai \cite{ajtai}, was that
for all $w$, there exists $n$ such that query time is
$\Omega(\sqrt{\lg w})$. Later, another bound was proved by Miltersen
\cite{milt}, who rephrased the same proof ideas in terms of
communication complexity: for all $n$, there exists $w$ such that
query time is $\Omega(\sqrt[3]{\lg n})$. Subsequently, the proof idea
was distilled even further into the technique of round elimination,
which was used to reprove the same bounds in a very concise way
\cite{mnsw}.
Beame and Fich \cite{bnf} proved two stronger bounds: for all $w$,
there exists $n$ such that query time is $\Omega(\lg w / \lg \lg w)$,
and for all $n$, there exists $w$ such that query time is
$\Omega(\sqrt{\lg n / \lg \lg n})$. They also gave a data structure
achieving $O(\min\{\frac{\lg w}{\lg \lg w}, \sqrt{\frac{\lg n}{\lg \lg
n}})\}$, which shows that these bounds are optimal if we insist on
pure bound in $n$ or $w$. These same bounds were also proved earlier
and independently by Xiao \cite{xiao}.
Beame and Fich's proof extends that of Ajtai, and is somewhat
complicated. Sen \cite{sen} later gave a stronger version of the round
elimination lemma, which gives a cleaner proof of the same bounds. In
this lecture, we will see how to use round elimination to prove the
predecessor lower bound, and we will sketch a proof of the round
elimination lemma.
The results we have discussed are actually for an easier problem: {\it
colored predecessor}. Each element is colored red or blue; a query on
an element $x$ returns the color of $x$'s predecessor. Because we can
solve colored predecessor using predecessor, a lower bound for colored
predecessor will yield a lower bound for predecessor. Having a lower
bound for the simpler problem is useful in reductions to other
problems.
\section{Communication Complexity}
We consider the problem in the communication complexity model. Let
Alice represent the query algorithm and Bob represent memory. Alice
has an input $x$, the query, and Bob has an input $y$, the contents of
the data structure. Alice and Bob are only permitted to communicate by
sending messages to each other of size at most $a$ and $b$
respectively; we will assume $a = O(\lg n)$, so it is possible to
address the entire data structure (using our polynomial space
assumption), and $b = w$, so a word of memory is returned. The goal is
to compute some function $f(x,y)$; in our case, the function is the
color of the predecessor. The parameter of interest is the number of
messages sent between Alice and Bob. This is at most twice the number
of probes needed in the cell-probe model. Note, however, that the
communication model is much stronger, since it allows both parties to
perform arbitrary computation (the memory can ``think'').
\subsection{The Predecessor Lower Bound}
We will prove an $\Omega(\min\{\lg_a w, \lg_b n\})$ lower bound on the
number of messages needed in the communication game. From this, we can
derive the Beame and Fich bounds. We have $a = \Theta(\lg n)$,
i.e.~the memory used is $n^{O(1)}$. Also, $b = w$. The lower bound is
worst (smallest) when
%
\[ \lg_a w = \lg_b n \Rightarrow
\frac{\lg w}{\lg \lg n} = \frac{\lg n}{\lg w} \Rightarrow
\lg^2 w = \lg n \lg \lg n \]
In terms of $n$, we find $\lg w = \sqrt{\lg n \lg \lg n}$, so the
bound becomes $\lg_a w = \sqrt{\frac{\lg n}{\lg \lg n}}$. In terms of
$w$, we find $\lg \lg w = \Theta(\lg \lg n)$, so the bound is $\lg_b n
= \frac{\lg w}{\lg \lg w}$.
\section{Round Elimination}
Round elimination can be applied to an abstract communication game
(not neccessarily related to the predecessor problem). It gives some
conditions under which the first round of communication can be
eliminated. To do this, we consider the ``$k$-fold'' of an arbitrary
function $f$:
\begin{definition}
Let $f^{(k)}$ be a variation on $f$, in which Alice has the $k$ inputs
$x_1, \ldots, x_k$, and Bob has inputs: $y$, $i \in {1, \ldots k}$,
and $x_1, \ldots, x_{i-1}$ (note that this overlaps with Alice's
inputs). The goal is to compute $f(x_i, y)$.
\end{definition}
Now assume Alice must send the first message. Observe that she must
send this message even though she doesn't know $i$ yet. Intuitively,
if $a \ll k$, with high probability she is unlikely to send anything
useful about $x_i$, which is the only part of her input that
matters. Thus, we can treat the communication protocol as starting
from the second message, eliminating the first.
\begin{lemma}[round elimination lemma]
Assume there is a protocol for $f^{(k)}$ where Alice speaks first that
uses $t$ messages and has error probability $\delta$. Then there is a
protocol for $f$ where Bob speaks first that uses $t-1$ messages and
has error probability $\delta + O(\sqrt{a/k})$.
\end{lemma}
One can give a good intuition for the result of this lemma. If $i$ is
chosen uniformly at random (which is the worst case), in Alice's first
message the expected number of bits ``about $x_i$'' that are sent is
$\frac{a}{k}$. Bob can guess these bits at random; the probability of
guessing all bits correctly is $2^{-a/k}$, so the probability of
failure is $1-2^{-a/k}$. The protocol should make an error either when
it did originally (with probability $\delta$), or when Bob guesses
these bits incorrectly. Because we are interested in small
$\frac{a}{k}$, we have $1 - 2^{-a/k} \approx a/k$. Thus, by
eliminating Alice's message, the error probability should increase by
about $\frac{a}{k}$. In the real world, this intuition is not entirely
correct, and we can only bound the increase in the error by
$\sqrt{a/k}$, but this is often enough for applications.
When proving a lower bound, the usual strategy is to use round
elimination repeatedly. We first eliminate Alice's first message.
Then, we interchange Alice and Bob in the lemma, and we eliminate
Bob's first message. We continue doing this, until we are left with no
message. If at this point we have a protocol with probability of
correctness bounded away from $\frac{1}{2}$, we usually have a
contradiction. This is because nontrivial functions of both intputs
cannot be computed without any communication, so the best possible
error probability is $\frac{1}{2}$ (random guessing). Observe that
each time we eliminate a message, we are changing the problem (we had
a protocol for $f^{(k)}$, and we obtain a protocol for $f$). This
means that inputs are getting smaller, and it limits the number of
times we can eliminate a round.
\section{Proof of Predecessor Bound}
Let $t$ be the number of cell probes (equivalently, the number of
rounds of communication) made by the predecessor algorithm. Our goal
is to perform $t$ round eliminations, leaving 0 messages. As we
perform more eliminations, we are reducing $n$ and $w$ to some $n'$
and $w'$. We want to increase the probability of error by at most
$\frac{1}{3t}$ each time, so that at the end, we still have a
nontrivial success probability (at least $\frac{2}{3}$). If, say, half
the elements are red and half are blue, the color of the predecessor
cannot be decided with probability better that $\frac{1}{2}$ (random
guessing) given no communication at all. So we reach a
contradiction. Our lower bound is the number of times we can do the
round elimination.
\subsection{Eliminating Alice-to-Bob}
Alice's input has $w'$ bits (initially, $w' = w$). Divide it into $k$
equal-size chunks $x_1, \ldots, x_k$, where $k = \Theta(at^2)$. Each
chunk is an integer of $w' / k$ bits.
We can construct a tree with branching factor $2^{w'/k}$ on the
$w'$-bit strings corresponding to the Alice's possible inputs, which
are the elements of the data structure. The tree then has height
$k$. This technique is reminiscent of van Emde Boas, in which we
divided the query word into two parts, and in $O(1)$ steps decided
that only one what interesting. For the lower bound, we need to divide
into more parts, but the idea is the same.
Because we are proving a lower bound, we are free to choose the inputs
to make the problem hard. Make the elements in the data structure have
a shared prefix of length $i$ with Alice's query and all differ in the
$i$-th chunk. Thus, all elements branch off from the query in the same
chunk (the hardest case for van Emde Boas). Alice and Bob know the
structure of the inputs, so Bob only needs to know $i$, the value of
the $i$-th chunk, and $x_1, \dots, x_i$ (because all of Bob's values
must start with this common prefix). Thus, when Alice's message is
eliminated, $w'$ is reduced to $w'/k = \Theta(w'/at^2)$. Using the
lemma, the error probability increases by $O(1/t)$, which is exactly
what we can afford per elimination.
\subsection{Eliminating Bob-to-Alice}
Now that Alice's message is eliminated, Bob is speaking first, so he
doesn't know the query value. Bob's input is $n'$ integers of $w'$
bits. Divide the set into $k$ equal chunks of $n'/k$ integers each,
where $k = \Theta(bt^2)$. Remember that fusion trees could recurse in
a set of size $n / w^{1/5}$ after $O(1)$ cell probes. Here, we are
proving that after one probe, you can only recurse into a set of size
$n / w^{O(1)}$, which gives the same bound (because the branching
factor is in the logarithm).
Again, we can construct a hard instance. We prefix the integers in
each chunk by a value of $\lg k$ bits, giving a unique indentifier of
each chunk (the $i$-th chunk starts with a prefix of $i$). Alice's
query starts with some random $\lg k$ bits, which decides which chunk
is interesting. If Bob speaks first, he cannot know which chunk is
interesting, so using the lemma, we can eliminate Bob's message the
error probability rises by $O(1/t)$.
The elimination reduces $n'$ to $n'/k = \Theta(n'/bt^2)$ and $w'$ to
$w' - \lg k = w' - \Theta(\lg bt^2)$. At long as $w'$ does not get too
small, $w = \Omega(\lg(bt^2))$, this last term is negligible (say, it
reduces $w'$ by a factor of at most $2$).
\subsection{Stopping}
Thus, each round elimination reduces $n'$ to $\Theta(n'/bt^2)$ and
$w'$ to $\Theta(w'/at^2)$. Further, the probability of error at the
end can be made to be at most $\frac{1}{3}$ by choosing appropriate
constants.
We stop the elimination when $w' = O(\lg(bt^2))$ or $n' = 2$. If these
stopping conditions are met, we have proven our lower bound: there
were many rounds initially, so we could do enough eliminations to
reduce $n$ and $w$ to these small values. Otherwise, we have a
protocol which gives an answer with zero messages, and the error
probability is at most $\frac{1}{3}$, which is impossible. So we must
be in the first case (the stopping conditions are met).
Hence, we established a lower bound $t = \Omega(\min\{\lg_{at^2} w,
\lg_{bt^2} n\})$. However, because $t = O(\lg n), a \ge \lg n$ and $t
= O(\lg w), b = w$, the bases of the logarithms are between $a$ and
$a^3$ and between $b$ and $b^3$ respectively. Thus, we found $t =
\Omega(\min\{\lg_a w, \lg_b n\})$.
\section{Sketch of the Proof for the Round Elimination Lemma}
\subsection{Some Information Theory Basics}
\begin{definition}
$H(x)$, called the entropy of $x$, is the number of bits needed on
average to represent a sample from a distribution of the random
variable $x$. Formally,
%
\[ H(x) = \sum_{x_0} \Pr[x=x_0] \cdot \lg \frac{1}{\Pr[x=x_0]} \]
\end{definition}
\begin{definition}
$H(x \mid y)$ is the conditional entropy of $x$ given $y$: the entropy
of $x$, if $y$ is known:
\[ H(x \mid y) = E_{y_0}[H(x|y=y_0)] \]
\end{definition}
\begin{definition}
$I(x:y)$ is the shared information between $x$ and $y$:
\[ I(x:y) = H(x) + H(y) - H((x,y)) = H(x) - H(x\mid y) \]
\end{definition}
$I(x:y \mid z)$ is defined in a manner similar to that of $H(x \mid
y)$.
\subsection{The Round Elimination Lemma}
Call Alice's first message $m = m(x_1,\ldots,x_k)$. Next, we use a
neat theorem from information theory to rewrite entropy as a sum:
%
\[ a = |m| \ge H(m) = \sum_{i=1}^k I(x_i:m \mid x_1,\ldots,x_{i-1}) \]
If $i$ is distributed uniformly in $\{1,\ldots,k\}$, then $E_i[I(x:m
\mid x_1,\ldots,x_k)] = \frac{H(m)}{k} \le \frac{a}{k}$. This is why
$\frac{a}{k}$ was an estimate for how many bits of information Bob
could learn from the message about Alice's message. Note that we
bounded $I(x_i : m \mid x_1, \ldots, x_{i-1})$, so even if Bob already
knows $x_1,\ldots,x_{i-1}$ and receives $m$, he still learns at most
$\frac{a}{k}$ bits about $x_i$.
The prove the lemma, we must build a protocol for $f$ given the
assumed protocol for $f^{(k)}$. We can build a protocol $f(x,y)$ as
follows:
\begin{enumerate}
\item Fix $x_1,\ldots,x_{i-1}$ and $i$ a priori (known to both
players) at random.
\item Alice pretends $x_i = x$.
\item Run the $f^{(k)}$ protocol, starting at the second message, by
assuming the first message is $m = m(x_1, \ldots, x_{i-1},
\tilde{x}_i, \ldots, \tilde{x}_k)$, where $\tilde{x}_j$ is a random
variable drawn from the distribution of $x_j$. Now the first message
does not depend on $x_i = x$ (even $x_i$ is chosen randomly), so Bob
can generate it by himself, without any initial message from Alice.
\item Now Alice has some actual $x$, which she must use as $x_i$, and
almost certainly $\tilde{x}_i \ne x$. But we know that $I(x_i:m)$ is
very small, so the message doesn't really depend on $x_i$ in a
crucial way. This means that a random message was probably good:
Alice can now fix $x_{i+1}, \dots, x_k$, so that $m(x_1, \ldots,
x_{i-1}, \tilde{x}_i, \ldots, \tilde{x}_k) = m(x_1, \ldots, x_{i-1},
x, \ldots, x_k)$, for the desired $x_i = x$.
\end{enumerate}
The last step is the crucial one, and it is what introduces an error
probability of $O(\sqrt{a/k})$. This is proved based on the ``Average
Encoding Theorem'' from information theory. There is also a more
subtle problem that this theorem solves: not only must $x_{i+1},
\dots, x_k$ exist, so that a Bob's random guess for a message is made
valid, but their distributions are close the the original
distributions, so the error probability $\delta$ does not increase too
much.
\bibliographystyle{alpha}
\begin{thebibliography}{78}
\bibitem[Ajt88]{ajtai} M. Ajtai: \emph{A lower bound for finding
predecessors in Yao's cell probe model}, Combinatorica 8(3): 235-247,
1988.
\bibitem[BF99]{bnf} P. Beame, F. Fich: \emph{Optimal Bounds for the
Predecessor Problem}, Symposium on the Theory of Computing 1999:
295-304.
\bibitem[Mil94]{milt} P. Miltersen: \emph{Lower bounds for
union-split-find related problems on random access machines},
Symposium on the Theory of Computing 1994: 625-634.
\bibitem[MNSW95]{mnsw} P. Miltersen, N. Nisan, S. Safra,
A. Wigderson: \emph{On data structures and asymmetric communication
complexity}, Symposium on the Theory of Computing 1995: 103-111.
\bibitem[Sen03]{sen} P. Sen: \emph{Lower bounds for predecessor
searching in the cell probe model}, IEEE Conference on Computational
Complexity 2003, 73-83.
\bibitem[Xia92]{xiao} B. Xiao: \emph{New bounds in cell probe model}, PhD
thesis, University of California, San Diego, 1992.
\end{thebibliography}
\end{document}