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\begin{document}

\lecture{4: Consistent Hashing, Graphs}{Jan 30,
  2020}{Instructor:\hspace{0.08cm}\emph{Alex Andoni }}{\emph{Arjav Nag (an2934), Yogesh Singh (ys3276)}}

\section{Background}

The original motivation behind Consistent Hashing arose from the need for web caching. Consider a fixed universe [$U$] of items (files in this case). The size of the universe would be very large $\approx 2^{(\text{length of file})}$.
\\\\
Given some subset of files $S\subset U$, with $|S|=m$. We want to cache these $m$ items in $n$ servers (buckets), where $n \ll m$. It is expected that the load on each server will roughly be $\approx \frac{m}{n}$. The assignment of files to servers can be done by a hash function $h: [U] \to [n]$.
\\\\
\noindent\textbf{Goal: }We want a mechanism to efficiently distribute the items when $n$ is increasing (adding new servers to the system).
\\\\
\noindent\textbf{Problem:} If we use hashing as a solution, with a hash function $h_n$, for a particular item $x$, we have:
$$h_n(x) = i$$
On adding a new server, a new hash function $h_{n+1}$ is used which could be randomized or from the family of hash functions. Then,
$$h_{n+1}(x) \neq i$$
with a probability $\approx 1-\frac{1}{n+1}$.
\\\\
This implies that most of the items would be rehashed and moved whenever a new server is added. This is inefficient. When adding server $n+1$, ideally we only want to move the items that map to it.








\section{Consistent Hashing}

Consider a random hash function for the items, $$h: [U] \to [M] \text{, where }M \gg m,n.$$
Also, consider a second random hash function for the servers, $$g: [N] \to [M]$$ where $N$ is an upper bound on the number of servers $n$. \\\\
$[M]$ can be thought of as an approximation for mapping to an interval of real numbers in $[0,1]$, since computers cannot represent reals precisely.

\begin{center}
\includegraphics[scale=0.5]{fig1.png} \\
\textbf{Figure 1: Circular representation for Consistent Hashing} \\
\textit{(Stanford CS168: The Modern Algorithmic Toolbox, Lecture 1: Introduction and Consistent Hashing, Tim Roughgarden \& Gregory Valiant)}
\end{center}

We can visualize a circle with each position representing a bucket in $[M]$. The set of items $S = \{x_0, x_1, x_2, ..., x_{m-1}\}$ and the set of servers $T = \{s_0, s_1, s_2, ..., s_{n-1}\}$ are all hashed to the corresponding buckets.
\\\\
The idea is to assign each object to the first server on its right. Thus, a server $s_i$ would be assigned the set of items $L_i$ between its location and the location of the preceding server on the circular representation:
$$L_i = \{x \in S, \quad s.t. \ g(s_j) < h(x) \leq g(s_i) \}$$
where $g(s_j),\ j \in [n] \text{ is maximal } g(s_j)$ to the left of $g(s_i)$.
\\

\begin{claim}
When a new server is added, we only need to move the items which it is servicing.
\end{claim}
\begin{proof}
In our initial assignment (Figure 1), we see that items $x_1$ and $x_2$ are serviced by server $s_0$. Now we introduce a new server $s_3$ (Figure 2). As per the defined mechanism, the first server on the right of $x_2$ is now $s_3$. For all other items, there is no change in this property and the servers servicing them remain unchanged. Hence, the only reassignment that needs to be done is to item $x_2$ which will be serviced by the new server. So only $x_2$ is moved from server $s_0$ to $s_3$.
\end{proof}

\begin{center}
\includegraphics[scale=0.65]{fig2.png} \\
\textbf{Figure 2: Circular representation after adding new server \bm{$s_3$}} \\
\textit{(Stanford CS168: The Modern Algorithmic Toolbox, Lecture 1: Introduction and Consistent Hashing, Tim Roughgarden \& Gregory Valiant)}
\end{center}


\begin{claim}
For $n$ servers, the expected load on a server $s_i$ is given by:
$$\bbE[|L_i|] = \frac{m}{n}$$
\end{claim}
\begin{proof}
This can be proved by the following two observations:
\begin{enumerate}
    \item By symmetry,
    $$\bbE[|L_i|] = \bbE[|L_j|] \quad \forall\ i,j \in [n]$$
    \item By definition,
    \begin{equation*}
        \begin{split}
            m & = \bbE\biggm[\sum_{i=0}^{n-1}|L_i|\biggm] \\
            & = n \cdot \bbE[|L_i|] \quad \forall\ i \in [n] \quad\text{ (from observation 1)}
        \end{split}
    \end{equation*}
\end{enumerate}

Thus,
$$\implies \bbE[|L_i|] = \frac{m}{n}$$
\end{proof}


\noindent(\textit{Note: }In the original hashing scheme, when
$\frac{m}{n}$ is sufficiently large (at least logarithmic), it can be
proved that the number of items in each server is $\approx
\frac{m}{n}$, up to a factor close to $1$, with high probability.) \\

\noindent While the expected load of each server is $\frac{m}{n}$, it is possible that when $n$ increases to $n+1$, there is one server whose load drops by a factor of 2 roughly. This is a relatively large variance.


\subsection{Analysis of Concentration}
In order to understand how well the algorithm concentrates items onto the servers, we need to compute the variance of the server load $l_i \triangleq |L_i|$. Since all servers are equivalent, by symmetry, we compute this for load $l_1$.
\begin{equation*}
    \begin{split}
        \var[l_1] & = \bbE[l_1^2] - (\bbE[l_1])^2 \\
        & = \bbE[l_1^2] - (\frac{m}{n})^2
    \end{split}
\end{equation*}
Now,
\begin{align*}
        \bbE_{h,g} [l_1^2] & = \bbE_g \biggm[ \bbE_h \biggm[\biggm( \sum_{x \in S} \bbI[x \in L_1] \biggm)^2\biggm]\biggm] \\
        & = \bbE_g \biggm[ \sum_{x,y \in S} \biggm( \underbrace{\bbE_h \biggm[\bbI[x \in L_1] \cdot \bbI[y \in L_1] \biggm]}_{\text{depends on width of interval between servers}}\biggm)\biggm]
\end{align*}
We define
the random variable $w$ to represent the distance to the server to the
left of server 1. Then,
\begin{align*}
        \bbE_{h,w} [l_1^2] & = \bbE_w \biggm[ \sum_{x,y \in S} \biggm( \bbE_h \biggm[\bbI[x \in L_1] \cdot \bbI[y \in L_1] \biggm]\biggm)\biggm] \\
        & = \bbE_w \biggm[ m\frac{w}{M} + m(m-1)(\frac{w}{M})^2 \biggm] \\
        & \leq \frac{m}{n} + \bbE_w \biggm[\frac{m^2w^2}{M^2 }\biggm]
\end{align*}
The second line in the above equation is obtained by grouping the expectations in the summation in two parts. The first group consists of expectations in which $x$ and $y$ are same. There are m such expectations(one for each item in $S$) and their value is $\frac{w}{M}$. The remaining expectations with different $x$ and $y$ are placed in the second group. Since distinct $x$ and $y$ are independent, these expectations are equal to $(\frac{w}{M})^2$.\\
$\bbE_w[ m\frac{w}{M}]$ is the expectation of the total number of items in server $s_1$. This can be thought of as each of the m items being assigned to server $s_1$ with probability $\frac{w}{M}$. We know this to be $\frac{m}{n}$. \\\\
$\bbE_w[\frac{m^2w^2}{M^2}]$ still needs to be computed:
$$
\bbE_w[\frac{w^2}{M^2}]=\int_{x\in [0,1]} \Pr_w[x<(w/M)^2]=\int_{x\in
  [0,1]} \Pr_w[w/M>\sqrt{x}]=\int_{x\in [0,1]} (1-\sqrt{x})^{m-1},
$$
as that's the probability that all the other $n-1$ servers do {\em not} fall
in the $w$ interval preceding the server 1. Hence, for $y=\sqrt{x}$, we have that
$$
\bbE_w[\frac{w^2}{M^2}]=\int_{y\in [0,1]} 2(1-y)^{n-1}y\le O(1/n^2)
$$

Hence we obtain:
\begin{claim}
$$\bbE[l_i^2] \leq O((\frac{m}{n})^2)$$
\end{claim}
% 
% MATERIAL TO BE ADDED
% 
% \begin{proof}
% \end{proof}


\noindent Applying Chebyshev's bound, with constant probability, we have:
\begin{align*}
    l_1 & = \frac{m}{n} \pm O(\sqrt{\var(l_1)}) \\
    & = \frac{m}{n} \pm O(\frac{m}{n})
\end{align*}
We have same problem as before. The variance is roughly proportional to the expectation. So there is significant deviation and the load for each server can vary by a constant factor. It can also not guarantee that a server will not have empty load.



\section{Variance Reduced Scheme}
We fix a parameter $k$ to control the reduction in variance (similar to the Morris+ algorithm). Instead of using the previous hash function $g$, we define a new function:
$$g: [N] \times [k] \to [M]$$
This can be visualized as each server having $k$ copies. Any server $s_i$ services items hashed between each of $g(s_{i,1}),g(s_{i,2}),...,g(s_{i,k})$ and the positions of the servers preceding them.

\begin{center}
\includegraphics[scale=0.55]{fig3.png} \\
\textbf{Figure 3: Representation of virtual server copies with \bm{$k=4$}} \\
\textit{(Stanford CS168: The Modern Algorithmic Toolbox, Lecture 1: Introduction and Consistent Hashing, Tim Roughgarden \& Gregory Valiant)}
\end{center}
This method increases the complexity when a new server is added. We will have to transfer the desired items to all the copies of the new server from the existing servers. In worst case, we might have to transfer items from $k$ different servers to the $k$ copies of the new server. However,  this also reduces the variance in the size of servers. This is because the items for the new server are fetched from various servers(instead of a singe) and hence the size of any existing server doesn't reduce drastically(as was the case previously). 
Let $vr\_load$ be the load on a server in the variance reduced scheme.
\begin{claim}
$$var(vr\_load) \leq O(var/k)$$
\end{claim}

If $l_1^{vr}$ is the load on server 1 in the variance reduced scheme, we can write
\begin{align*}
    l_1^{vr} & = \frac{m}{n} \pm O(\frac{m}{n}).\frac{1}{\sqrt{k}}
\end{align*}

\section{Introduction to Graphs}
A graph can be defined as a set of vertices(V) and edges(E) between the vertices where each edge can have an associated weight/length(W) and a capacity(C).

$$G=(V, E, W:E\to R, C:E\to R_+)$$
\\
There are some basic graph algorithms like breadth-first search, Dijkstra's algorithm which are used to find shortest paths. These algorithms are good to know but won't be covered in this class. Such problems can be solved in linear time in terms of number of edges and nodes.\\
\\
We will be focusing on algorithms like maximum flow which can't be solved in linear time. In the next lecture we will define the maximum flow problem and talk about a basic solution using the Ford-Fulkerson algorithms. Then we will discuss how to solve it faster than Ford-Fulkerson. 

\end{document}
