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\hbox to 5.78in { {\bf Sketching Algorithms for Big Data } \hfill #2 }
\vspace{4mm}
\hbox to 5.78in { {\Large \hfill #5 \hfill} }
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\begin{filecontents}{mybib.bib}
@inproceedings{alonFOCS,
author = {Noga Alon and
Bo'az Klartag},
title = {Optimal Compression of Approximate Inner Products and Dimension Reduction},
booktitle = {58th {IEEE} Annual Symposium on Foundations of Computer Science, {FOCS}
2017, Berkeley, CA, USA, October 15-17, 2017},
pages = {639--650},
year = {2017},
crossref = {DBLP:conf/focs/2017},
url = {https://doi.org/10.1109/FOCS.2017.65},
doi = {10.1109/FOCS.2017.65},
timestamp = {Thu, 16 Nov 2017 15:09:35 +0100},
biburl = {http://dblp.org/rec/bib/conf/focs/AlonK17},
bibsource = {dblp computer science bibliography, http://dblp.org}
}
@inproceedings{Kushilevitz,
author = {Eyal Kushilevitz and
Rafail Ostrovsky and
Yuval Rabani},
title = {Efficient Search for Approximate Nearest Neighbor in High Dimensional
Spaces},
booktitle = {Proceedings of the Thirtieth Annual {ACM} Symposium on the Theory
of Computing, Dallas, Texas, USA, May 23-26, 1998},
pages = {614--623},
year = {1998},
crossref = {DBLP:conf/stoc/1998},
url = {http://doi.acm.org/10.1145/276698.276877},
doi = {10.1145/276698.276877},
timestamp = {Thu, 16 Feb 2012 12:06:54 +0100},
biburl = {http://dblp.org/rec/bib/conf/stoc/KushilevitzOR98},
bibsource = {dblp computer science bibliography, http://dblp.org}
}
@inproceedings{indyk,
author = {Piotr Indyk and
Tal Wagner},
title = {Near-Optimal (Euclidean) Metric Compression},
booktitle = {Proceedings of the Twenty-Eighth Annual {ACM-SIAM} Symposium on Discrete
Algorithms, {SODA} 2017, Barcelona, Spain, Hotel Porta Fira, January
16-19},
pages = {710--723},
year = {2017},
crossref = {DBLP:conf/soda/2017},
url = {https://doi.org/10.1137/1.9781611974782.45},
doi = {10.1137/1.9781611974782.45},
timestamp = {Wed, 24 May 2017 08:31:21 +0200},
biburl = {http://dblp.org/rec/bib/conf/soda/IndykW17},
bibsource = {dblp computer science bibliography, http://dblp.org}
}
@inproceedings{larsen,
author = {Kasper Green Larsen and
Jelani Nelson},
title = {Optimality of the Johnson-Lindenstrauss Lemma},
booktitle = {58th {IEEE} Annual Symposium on Foundations of Computer Science, {FOCS}
2017, Berkeley, CA, USA, October 15-17, 2017},
pages = {633--638},
year = {2017},
crossref = {DBLP:conf/focs/2017},
url = {https://doi.org/10.1109/FOCS.2017.64},
doi = {10.1109/FOCS.2017.64},
timestamp = {Thu, 16 Nov 2017 15:09:35 +0100},
biburl = {http://dblp.org/rec/bib/conf/focs/LarsenN17},
bibsource = {dblp computer science bibliography, http://dblp.org}
}
\end{filecontents}
\newcommand{\atime}[1]{\texttt{[#1]}}
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\begin{document}
\lecture{23 --- November 21, 2017}{Fall 2017}{Prof.\ Noga Alon}{Preetum Nakkiran}
\section{Overview}
This lecture will be on
``Optimal Compression of Approximate Inner Products and
Dimension Reduction'',
following the paper of Alon and Klartag~\cite{alonFOCS}.
We will study the following object:
\begin{definition}
An \emph{$\eps$-inner-product sketch}
for points $X \subseteq \R^k$, $|X| = n$, such that $\forall x: ||x||_2 \leq 1$,
is a data structure that enables one to compute
the inner-product
$$\innp{x, x'} ~\forall x, x' \in X$$
up to additive error $\eps$.
\end{definition}
Let $f(n, k, \eps)$ be the minimum size (in bits) of an
$\eps$-inner-product sketch for $n$ points in dimension $k$.
(We can assume $n \geq k$ WLOG, and further $\eps \geq 1/n^{0.5 - \delta}$,
since this is the interesting case for dimensionality reduction).
There are various applications for this data structure,
for example in streaming, compressed sensing, and computational geometry (eg,
Nearest-Neighbors).
{\bf Main Questions:}
What is $f(n, k, \eps)$? And, can we compute sketches efficiently?
We will first do a warm-up, and then present the Main Theorem of~\cite{alonFOCS}
which exactly characterizes $f(n, k, \eps)$.
\section{Warm-Up}
First, let us consider the basic case of when $k = n$.
We claim an easy upper-bound here is
$$f(n, n, \eps) \leq O(n \frac{\log n}{\eps^2} \log(1/\eps))$$
To do this:
\begin{enumerate}
\item First use JL to project the $n$ points in $\R^n$ down to dimension
$m := O(\frac{\log n}{\eps^2})$, with distortion $O(\eps)$.
\item Then, round the projected points to an $\eps$-net of the unit ball in $\R^k$.
\end{enumerate}
The $\eps$-net in Step 2 requires at most $O(\frac{1}{\eps})^{(1+o(1))m}$ points,
so we need $O(m \log(1/\eps))$ bits per point.
This yields the claim.
In fact, this bound on $f(n, n, \eps)$ is nearly-tight.
The $\log(1/\eps)$ factor can be shaved, as first shown by
Kushilevitz, Ostrovsky, and Rabani~\cite{Kushilevitz}.
The construction is simple:
{\bf Construction:} Let $v_1, \dots, v_t$ be random unit vectors in $\R^n$, for
$t = O(\frac{\log n}{\eps^2})$.
For every $x \in X$, maintain just the $t$ signs:
$\{\text{sign}(\innp{x, v_i})\}_{i \in [t]}$.
{\bf Proof sketch:}
The idea is, for two vectors $x, x'$, a random unit vector $v_i$ will
distinguish them (ie, $sign(\innp{x, v_i}) \neq sign(\innp{x', v_i})$)
with probability proportional to the angle between $x, x'$.
Thus, maintaining $t$ signs gives a good enough approximation of this angle,
and taking the cosine recovers approximately $\innp{x, x'}$.
\qed
Note that this construction does not yield better bounds on $f(n, k, \eps)$
for smaller $k < n$.
\section{Main Theorem}
\begin{theorem}[\cite{alonFOCS}]
Let $f(n, k, \eps)$ be the minimum size (in bits) of an
$\eps$-inner-product sketch for $n$ points in dimension $k$.
Then, $\forall n \geq k, \eps \geq \frac{1}{n^{0.49}}$:
\begin{enumerate}[(A)]
\item For $\frac{\log n}{\eps^2} \leq k \leq n:$
$f(n, k, \eps) = \Theta(\frac{n\log n}{\eps^2})$
\item For $\log n \leq k \leq \frac{\log n^2}{\eps^2}:$
$f(n, k, \eps) = \Theta(nk \log(2+\frac{\log n}{\eps^2k}))$
\item For $1 \leq k \leq \log n:$
$f(n, k, \eps) = \Theta(nk \log(1/\eps))$.
\end{enumerate}
\end{theorem}
Note that these bounds all agree on their common boundaries.
\begin{remark}
The Main Theorem considers \emph{additive} error.
Recently Indyk and Wagner ~\cite{indyk}
have shown how to achieve \emph{relative} error guarantees (which is harder).
Note, there is still a gap of $\log(1/\eps)$ between upper and lower bounds here.
\end{remark}
\begin{remark}
The Main Theorem gives an alternative proof of a recent result by Larsen and
Nelson~\cite{larsen}:
There does not exists a dimensionality reduction of $n$ points in $\R^n$
to dimension $< \frac{c\log n}{\eps^2}$ for some small constant $c$.
This follows because in the regime $k < \frac{\log n}{\eps^2}$,
$f(n, k, \eps)$ decays rapidly with $k$, in particular
$f(n, n, 2\eps) > f(n, k, \eps)$.
That is, there are simply too many configurations of points in $\R^n$
to be faithfully compressed in $\R^k$.
Note, this also implies that once $k < \frac{n}{\eps^2}$, even dimensionality
reduction by a factor of (say) $10$ is impossible.
\end{remark}
\begin{remark}
It turns out that the Main Theorem can be combined with the
\emph{Khatri-Sidak Lemma} and \emph{Hargi\'e Inequality}
(which are special cases of ``Gaussian Correlation''),
to work for all $\eps \geq \frac{1}{\sqrt{n}}$.
See~\cite{alonFOCS} for details.
For the statement, replace $\log n$ by $\log(2+\eps^2 n)$ everywhere in the
Main Theorem.
\end{remark}
\section{Proof of Main Theorem}
We will now see two proofs of the upper bound (non-constructive and
constructive), and one proof of the lower bound.
\subsection{Upper Bound}
Consider first the regime (A),
where $\frac{\log n}{\eps^2} \leq k \leq n$.
We wish to show $f(n, k, \eps) = \Theta(\frac{n\log n}{\eps^2})$.
\paragraph{Regime (A)}
By monotonicity, it is sufficient to show (A) for $k = n$.
By JL, we may first project points into dimension $m := \frac{\log n}{\eps^2}$,
with distortion $O(\eps)$.
For vectors $w_1, \dots, w_n \in \R^m$, define the \emph{Gram Matrix}
$G(w_1, \dots, w_n)$ as $G_{i, j} := \innp{w_i, w_j}$.
Say two Gram Matrices $G, G'$ are \emph{$\eps$-separated} if
$\exists i \neq j:~ |G_{i, j} - G'_{i, j}| > \eps$.
Let $\G$ be a maximal (w.r.t containment) set of $\eps$-separated Gram Matrices.
Then, we clearly have $f(n, k, \eps) \leq \log |\G|$
(simply by remembering, for a point set $X$,
the index of a $g \in \G$ that $G(X)$ is not $\eps$-separated from)
We now want to bound $|\G|$.
We will use essentially a volume argument.
Let $v_1, \dots, v_n \in \R^m$ be random vectors, each uniform in a ball of
radius $2$ about the origin.
For all $G(w_1, \dots w_n) \in \G$, define the event
$$A_G := \{ \forall i, j: |\innp{v_i, v_j} - \innp{w_i, w_j}| \leq \eps/2\}.$$
Notice the events $\{A_G\}_{G \in \G}$ are pairwise disjoint, by
$\eps$-separatedness of $\G$.
Thus, $\sum_{G \in \G} \Pr[A_G] \leq 1$.
So if the individual probabilities $\Pr[A_G]$ are large, then there cannot be
many such events $A_G$ (thus bounding $|\G|$).
It is easy to see
that $\Pr[A_G] \geq \Omega(\eps)^{mn}$
(by noticing that the condition
$\{\forall i: ||v_i - w_i||_2 \leq \eps/4\}$ is sufficient).
We will show that in fact, a much better bound holds:
$$\Pr[A_G] \geq \Omega(1)^{mn}.$$
First, let us condition on the event
$$E := \{\forall i: ||v_i - w_i||_2 \leq 1\}.$$
We have $\Pr[E] = (\frac{1}{2})^{nm}$, by comparing volumes (since we pick $w_i$
uniformly from a ball of radius $2$).
Moreover, conditioned on this event $E$, the vector $(v_i - w_i)$ is uniform
over a unit ball in $\R^m$. Thus, by concentration (eg Azuma-Hoeffding) we have
$$\forall i \neq j: \Pr[|\innp{v_i - w_i, w_j}| \geq \eps/4 ~~| E] \leq 2 e^{-\Omega(\eps^2 m)} < \frac{1}{2n^2}$$
And symmetrically,
$$\forall i \neq j: \Pr[|\innp{v_i, v_j - w_j}| \geq \eps/4 ~~| E] \leq 2 e^{-\Omega(\eps^2 m)} < \frac{1}{2n^2}$$
This allows us to union bound over all pairs $i \neq j$.
In particular, we now know that with probability $\geq \frac{1}{2}
(\frac{1}{2})^{nm}$, the following simultaneously holds:
$$
\begin{cases}
\forall i: ||v_i - w_i||_2 \leq 1\\
\forall i \neq j: |\innp{v_i - w_i, w_j}| < \eps/4\\
\forall i \neq j: |\innp{v_i, v_j - w_j}| < \eps/4
\end{cases}
$$
This is sufficient to imply $A_G$ holds, since by the last two conditions and
triangle inequality:
$$
|\innp{v_i, v_j} - \innp{w_i, w_j}|
=
|\innp{v_i - w_i, v_j} + \innp{w_i, v_j- w_j}|
\leq
|\innp{v_i - w_i, v_j}| + |\innp{w_i, v_j- w_j}|
\leq \eps/2
$$
Thus, $\Pr[A_G] \geq (\frac{1}{2})^{mn+1}$ and so $|\G| \leq 2^{mn +1}$.
And we have
$f(n, k, \eps) \leq \log |\G| = O(nm) = O(n \frac{\log n}{\eps^2})$ as desired.
\qed
\begin{remark}
What is going on is, once we condition on $E$, then
each $w_i$ is uniform in a unit ball centered around $v_i$,
and so and so by concentration
we have $\innp{w_i, w_j} \approx \innp{v_i, v_j}$.
\end{remark}
\paragraph{Regime (B)}
This is essentially the same as Regime (A), except without the JL projection.
In particular, for $k = \frac{\delta^2}{\eps^2}\log n$ and $2\eps \leq \delta < 1/2$:
Let $\G$ be a maximal set of $\eps$-separated Gram Matrices (of vectors in
$\R^k$ with norm $\leq 1$).
Take $v_1, \dots, v_n \in \R^k$ random vectors as before, and condition on the
event
$$E := \{\forall i: ||v_i - w_i||_2 \leq \delta/40 \}$$
Then,
$$\Pr[E] = (\delta/80)^{kn}$$
and the rest follows as before.
That is, after conditioning on $E$,
$$\forall i \neq j: \Pr[|\innp{v_i - w_i, w_j}| \geq \eps/4 ~~|E]
\leq 2 e^{-\Omega(k\eps^2/\delta^2)} < \frac{1}{2n^2}$$
And we conclude as in part (A).
\qed
\paragraph{Regime (C)}
Here we can simply round to the closest point in an $\eps$-net for the unit
ball. This gives $O(k \log(1/\eps))$ bits per point.
\qed
\subsection{Upper Bound: Algorithmic Proof}
Here we sketch an alternative, constructive proof of the upper-bound.
\paragraph{Regime (A)}
Apply JL projection into dimension $m := \frac{\log n}{\eps^2}$,
and then do \emph{randomized} rounding to a $(1/2)$-net.
(For example, each coordinate is randomly rounded to an integral multiple of
$\frac{1}{\sqrt{2m}}$, preserving its expectation).
By concentration (eg, Chernoff-Hoeffding) the randomized rounding
stage does not distort too much w.h.p.
\paragraph{Regime (B)}
Here we skip the JL projection, and just do randomized rounding
of each coordinate to an integral multiple of $\delta/\sqrt{k}$
for $k = \frac{\delta^2}{\eps^2}\log n$.
\paragraph{Regime (C)}
We simply (deterministically) round each coordinate to the nearest integer multiple of
$\eps/\sqrt{k}$.
\subsection{Lower Bound}
We now describe the main ideas of the lower bound.
The Main Lemma is the following:
\begin{lemma}
If $k = \frac{\delta^2}{200 \eps^2} \log n$,
then
$$f(n, k, \eps) \geq \Omega(n k \log(1/\delta)).$$
\end{lemma}
\paragraph{Proof Sketch}
Let $N$ be a maximal $\delta$-packing in $\R^k$ (such that the distance between any two
points in $N$ is $\geq \delta$).
Note $|N| = (\frac{1}{\delta})^k$.
We can show that w.h.p., a random set $R$ of $(n/2)$ unit vectors in $\R^k$
satisfies:
$$\forall x \neq x' \in N: \exists y \in R: |\innp{x, y} - \innp{x', y}| >
\eps$$
That is, w.h.p. $R$ distinguishes every pair of points in $N$.
Fix such an $R$. Then, each (ordered) configuration of
$R$ together with $> (n/2)$ points of $N$ requires a \emph{distinct encoding}.
Because, any two such configurations must share two points $x, x' \in N$ in common,
and then $R$ suffices to distinguish these $x, x'$ -- so they cannot be
represented by the same encoding.
Thus,
$$f(n, k, \eps) \geq \log |N|^{n/2} = \Omega(\frac{n}{2} \log |N|) = \Omega( nk
\log(1/\delta))$$
\qed
\bibliography{mybib}
\bibliographystyle{alpha}
\end{document}