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\head{CS G399}{Algorithmic Power Tools I}{Fall 2007}{17 November
2007}{Abhishek Chaubey}


\newcommand{\p}{\mathcal{P}}
\newcommand{\U}{\mathcal{U}}
\newcommand{\coll}{\mathcal{S}}
\newcommand{\xpy}{x + \lambda y}
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\newcommand{\opt}{\mbox{OPT}}
\newcommand{\eps}{\varepsilon}

\noindent {\bf Lecture Outline:}
\begin{itemize}
\item
Max-Cut Recap
\item
Sparsest Cut
\begin{itemize}
\item
Metrics and Embeddings
\item
LP based $O(\log n)$
\item
LP based $O(\sqrt{\log n})$

\end{itemize}
\end{itemize}

In this lecture we will review the Max-Cut which we did in the last class.
Then we will study Sparsest-Cut along with notions of metrics, norms
and embedding.

\section{Max-Cut Review}

In the last class we saw an algorithm for Max-Cut using SDP as a
black box. In Max-Cut, we break the graph into two pieces, to
maximize the number of edges going across. The applications of
Max-cut are in fields of VLSI design where we may desire high
bandwidth/data-rates across two pieces.

We wrote vector program [P] as follows:

\[
 \max \frac{1}{2} \sum_{(i,j) \in E} {(1-v_i.v_j)}
\]
\[
 \mbox{such that } v_i \cdot v_i = 1
\]

The optimal value of [P] $\ge$ Max-Cut. Here, we are replacing the
naive upper bound by $Z_P^*$ and hence getting a better
approximation ratio. Cut is discrete and binary in nature i.e.
either the vertex will be in $S$ or in $\bar{S}$. So approach of the
algorithm is to make it continuous, solve it and then use that
solution to guide the discrete. For further analysis of this
algorithm, refer to last lecture.

Lovasz realized that algorithmic ideas from Linear Programming can
be extended to Semi-Definite Programming. But SDP had kind of
continuous relaxation and getting something discrete out of it was
challenging. The paper by Goemans-Williamson was the first
demonstration in which we can write the SDP and which can then be
rounded cleverly to get a good approximation ratio.


\section{Sparsest-Cut Problem}

The aim in Sparsest Cut is to break the graph into two pieces such
that the pieces are large but the number of edges crossing the cut
is small. In other words, we want small number edges connecting two
large pieces.

{\bf Problem:}

Given a graph $G = (V,E)$ and $w: E \rightarrow Z^+$. Determine cut
$(S,\bar{S})$ such that $\frac{E(S,\bar{S})}{|S||\bar{S}|}$ is
minimized where $E(S,\bar{S})$ is number edges crossing $S$ and
$\bar{S}$.

{\bf Motivation:}

Sparsest-Cut Problem is much better motivated as compared to Max-Cut
Problem. One motivation for Sparsest-Cut is Balanced-Cut (where we
have two pieces with small number of edges crossing and each piece
is constant fraction of $n$). Example: (1/3,2/3) balanced cuts,
$(\alpha,(1-\alpha))$ balanced cuts.

Balanced Cuts are widely studied because it is very fundamental to
Divide and Conquer paradigm. It is because if we can break our
problem instance into balanced pieces, then the algorithm would be
better in terms of running time and quality. From efficiency
standpoint if both the pieces are some constant fraction of total
size, it is guaranteed that we are making progress. From quality of
algorithm standpoint, if the  instance is broken evenly and the
number of edges crossing between is small, then when the solutions
of each pieces are stitched together, the loss of stitching the
constant fraction pieces is not too much. As Divide and Conquer is a
basic paradigm in algorithm design, people care about Balanced-Cut.
It can be shown that if we can get a approximation for Sparsest-Cut,
then we can get a pseudo-approximation to Balanced-Cut.

Another motivation for Sparsest-Cut is counting substructures in
graphs. For several counting applications, random walks on graphs
are required. And to do random walks on graph, it is important to
show that we get good counts if the random walks converge to a
stationary distribution. Intuition is to randomly walk on a graph
for a long time, and after sometime get lost in the graph (i.e.
after getting lost it is equally likely to be present anywhere in
the graph). And getting lost is a necessary requirement for getting
substructures (i.e. this means that you count substructures in
different parts of the graphs). The size of the sparsest cut is
related to the time taken for random walks to converge.

\subsection{Algorithm}

We discussed Leighton-Rao(LR) algorithm couple of weeks back which
is a combinatorial algorithm for uniform case. We will present
$O(\log n)$ due to Linial-London-Rabinovich (LLR) and
Auman-Rabani(AR) which trivially handles non-uniform case.

As per LLR, we take $\frac{E(S,\bar{S})}{|S||\bar{S}|}$ as
optimization over a $metric$.

\begin{Definition}
\label{def:metric} If $S$ is a set of points and $R^{\ge 0}$ is a
set of positive reals then {\bf metric} is a function $d: (S * S)
\rightarrow R^{\ge 0}$ satisfying:

i. $Strict Metrics: $ $d(S_i,S_j) =  0 \Leftrightarrow S_i =S_j$

ii. $Symmetry: $ $d(S_i,S_j) = d(S_j,S_i)$

iii. $Triangle Inequality: $ $d(S_i,S_j) + d(S_i,S_j) \ge
d(S_i,S_k)$
\end{Definition}

\begin{Definition}
\label{def:norm} Suppose we have $m$-dimensional vector $\vec{v_i}$
i.e. $\vec{v_i} \in R^m$, then {\bf norm} (represented by $\| \|$)
can be defined as a function $ \| \|: R^m \rightarrow R^{\ge 0}$. We
can think of a norm as the length of the vector. It satisfies the
following:

i. There is only vector whose norm is zero and that is zero vector
i.e. $\|\vec{v}\| = 0 \Leftrightarrow \vec{v}=\vec{0}$

ii. Linearity of a norm under scalar multiplication (if we scale a
vector by some amount, its norm also scales by the same amount) i.e.
$ \|\lambda\vec{v}\| = |\lambda|\|\vec{v}\|$

iii. Triangle Inequality i.e. $ \|\vec{u} + \vec{v}\| \le
\|\vec{u}\| + \|\vec{v}\|$
\end{Definition}

{\bf Examples of norm: }
\begin{itemize}
\item
$L_1 = \sum |v_i|$
\item
$L_2 = \sqrt{\sum |v_i|^2}$
\item
$L_n = \sqrt[n]{\sum |v_i|^n}$ where $n$ is a real number $\ge$ 1
\item
$L_{\inf} = \max_i |v_i|$
\end{itemize}

Consider $R^m$. Take any centrally symmetric convex ($CSC$) body.
Centrally Symmetric means if $\vec{u} \in body$ then $-\vec{u} \in
body$. $Convex$ means if $\vec{u}, \vec{v} \in body$ then $\alpha
\vec{u} + (1-\alpha)\vec{v} \in body$.

Now it is to be noted that every norm corresponds one-to-one to
$CSC$. In other words, $CSC$ induces a norm. To reduce a $CSC$ to a
norm, for any vector, we will compute the $\lambda$ by which, the
vector need to be scaled to bring it to the boundary of the ball.
Since $CSC$ is convex and centrally symmetric, there is exactly one
point at which the vector can be scaled to the boundary. Then,
$\lambda$ is the norm of the vector. This norm satisfies $linearity$
trivially. The $triangle$ $inequality$ comes from the $convexity$ of
CSC. The following are $CSC$ corresponding to some norms.

\begin{figure}[htb!]
  \begin{center}
  \includegraphics[height=60mm]{csc.png}
  \end{center}
\end{figure}


Norms induce metrics. Suppose we have bunch of vectors in $R^m$,
then pick any $CSC$ body, which induces a norm. And define metric as
$d(\vec{u},\vec{v}) = \|\vec{u}-\vec{v}\|$.

$l_1$ is used to denote family of metrics. In particular $l_1^m$
denotes that there are $n$ points belonging to $m$-dimensional space
and the distance between them is induced by the $L_1$ norm.

\begin{Definition}
\label{def:embedding} We take a metric $S$, with distance function
$d$ on it and embed into another metric $T$ with distance function
$d'$. $Embedding$ $(E)$ can be formally defined as

\[ E:(S,d) \ce{<=>[\text{$D$}][\text{$D$}]}  (T,d')
\]

If $i$ and $j$ are the two points, then ratio of $d(i,j)$ and
$d'(E(i),E(j))$ is captured by $distortion$ $(D_{i,j})$ for the edge
$(i,j)$. We will focus our attention on non-expansive embedding,
that is $D_{i,j} \le 1$. $Distortion$ $(D)$ for the entire
$embedding$ is $\max_{(i,j)} D_{i,j}$. The notion of $Average
Distortion$ can be defined as $\frac{\sum_{i,j} D_{i,j}}{\mbox
{number of pairs}}$.
\end{Definition}

Consider the following graph. Take a subset S.  Distance between any
pair of vertices on the same side is 0, else if they are on opposite
side then 1 (irrespective of if there exists an edge between them).
We claim that any $cut$ induces a $metric$ called $cut$ $metric$
$d_S$ (note that distance defined satisfies non negativity, strict
metrics, symmetry and triangular inequality).

\begin{figure}[htb!]
  \begin{center}
  \includegraphics[height=60mm]{cut_metric.png}
  \end{center}
\end{figure}

Recall, we said Sparsest Cut as optimization over a $metric$.
\begin{eqnarray}
\mbox{Sparsest Cut} &=& \min_S \frac{E(S,\bar{S})}{|S||\bar{S}|}\nonumber\\
                    &=& \min_{d_S} \frac {\sum_{i,j \in E} d_S(i,j)}{\sum_{i,j}
                        d_S(i,j)} \nonumber\\
                    &=& \min_{d_{l_1}} \frac {\sum_{i,j \in E} d_{l_1}(i,j)}{\sum_{i,j}
                        d_{l_1}(i,j)}
\end{eqnarray}

Note the equality (1). It says, we can relax from $cut$ $metrics$ to
$l_1$ $metrics$. We now argue the last step of the above equation.
It can be observed that any $cut$ $metric$ is an $l_1$ $metric$.

One claim is, if we take at any $l_1$ metric, it is equal to the
combination of $cut$ $metrics$ for some collection of cuts. In other
words, $l_1$ is $the$ cut cone. So, if we take any metric induced by
$manhattan$ $distance$ $(L_1)$, then that can be represented by sum
of $cut$ $metrics$. Also, the reverse direction is true i.e. any
point that is represented as some combination of cuts, from that we
can get a $l_1$ embedding. (Note that the reverse direction is not
required for our analysis). Thus, there exists scalar $\lambda_S$,
$S \subseteq V$ such that

\[
d_{l_1} = \sum_{S \subseteq V}(\lambda_S d_S)
\]

This claim doesn't imply equality (1) as an $l_1$ metric we have is
sum of $cut$ $metrics$ (not one particular cut). We can break the
equality (1) as follows:

\[
\min_{d_{l_1}} \frac {\sum_{i,j \in E} d_{l_1}(i,j)}{\sum_{i,j}
d_{l_1}(i,j)} = \min_{S_1 \ldots S_k} \frac {\sum_{i,j \in E}
\sum_{S_t} \lambda_{S_t}d_{S_t}(i,j)}{\sum_{i,j} \sum_{S_t}
\lambda_{S_t}d_{S_t}(i,j)}
\]

Using $\frac{\sum_i a_i}{\sum_i b_i} \ge min_i \frac{a_i}{b_i}$ and
the above equality (just switch the order of summation and move
$\lambda$ outside), we can argue that the equality (1) holds.

Now,we know that we can optimize over $l_1$. If we remove $\lambda$
from both numerator and denominator, the ratio remains the same.
Hence, we can write LP for sparsest cut as follows:

\[
 \min_{d_{l_1}} \sum_{i,j \in E} d_{l_1}(i,j)
\]
\[
 \mbox{such that } \sum_{i,j} d_{l_1}(i,j) =1
\]

Now, from $l_1$, we can relax it to any metric. After relaxation the
value of objective function will reduce. But we can add additional
constraint (triangle inequality) to represent it as LP.

\[
 \min \sum_{i,j \in E} d_{l}(i,j)
\]
\[
 \mbox{such that } \sum_{i,j} d_{l}(i,j) =1
\]
 $\forall i, j, k \,\, d_l(i,j)+d_l(j,k) \ge d_l(i,k)$

We will solve this LP, get a metric and use the following theorem by
Bourgain.

{\bf Bourgain:}

Any metric \ce{->[\text{$O(\log n)$}]} $l_1^{O(\log^2n)}$ i.e. any
metric embeds in $l_1$ with $O(\log n)$ distortion.

{\bf Roadmap to Solution:}

We wanted to minimize over cut metrics. Then we said that minimizing
over cut metrics is same as minimizing over $l_1$. Then we relaxed
it and minimize over any metric. By Bourgain,  any metric embeds in
$l_1$ with $O(\log n)$ distortion. From $l_1$, we can induce the cut
($l_1$ and cut cone are same).

The roadmap for Arora, Rao, Vazirani is similar. We want to minimize
over cut metrics. Then minimizing over cuts is same as minimizing
over $l_1$. Then we relaxed it and minimize over $l_2^2$ metric
(this metric is captured by an SDP). By ARV, $l_2^2$ embeds in $l_1$
with $O(\sqrt{\log n})$ average distortion. From $l_1$, we can
obtain the cut.

\end{document}
