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\begin{document}
\lecture{5 --- February 26, 2007}{Spring 2007}{Prof.\ Erik
Demaine}{Katherine Lai}
\section{Overview}
In the last lecture we discussed the link-cut tree: a dynamic tree
that achieves $O(\lg n)$ amortized time per operation. In this
lecture, we leave the BST model and enter the pointer-machine model.
In this model, we are allowed to have nodes which each hold $O(1)$
pointers and $O(1)$ fields (e.g. integers). Our machine is also
allowed to keep track of $O(1)$ ``fingers'' to nodes at a time. The
operations we are allowed to do are to copy a finger, follow a pointer
to a node, and to retrieve or change the contents of a node's fields.
So, the BST model allows for a subset of the data structures allowable
in the pointer-machine model. We will focus on data structures in
this model that solve the dynamic connectivity problem.
\section{Euler tour trees}
The Euler tour tree data structure is due to Henzinger and King in
\cite{henzinger}. An Euler tour of a graph is a path that traverses
each edge exactly once. In the context of a tree, we say that each
edge is bidirectional, so the Euler tour is the path along the tree
that begins at the root and ends at the root, traversing each edge
exactly twice --- once to enter the subtree at the other endpoint and
once to leave it. You can think of an Euler tour as just being a
depth first traversal where we return to the root at the end.
We will store the Euler tour of the tree as a balanced binary search
tree with one node for each time a node in the represented tree was
visited, with each node in the tree keyed by its time of visit. Each
node in the represented tree will also hold pointers to the nodes
representing the first and last time it was visited.
\begin{figure}[htp]
\centering
\includegraphics[totalheight=0.3\textheight]{euler.jpg}
\caption{We represent a tree's Euler tour as a sequence of visitations
starting and ending at the root. Pointers from that sequence point
to the node, and pointers from the node point to the first and last
visitations. (Only the pointers relating to the root are shown.)
The sequence of visitations is then stored in a balanced binary tree
(not shown).}
\end{figure}
While the link-cut trees we discussed last lecture are good for
maintaining aggregates on paths of a tree (making it a good choice
data structure in network flow algorithms), Euler tour trees are
better at keeping aggregate information on subtrees. We will use this
feature of Euler tour trees toward the end of the lecture notes.
We want Euler tour trees to be able to perform the same three
operations as link-cut trees:
\begin{itemize}
\item[-] \textsc{{\textbf Findroot}}\textbf{(v)}\\
In the Euler tour, the root will be visited first (and last), so we
simply have to find the min (or max) in the balanced BST.
\item[-] \textsc{{\textbf Cut}}\textbf{(v)}\\
The Euler tour of $v$'s subtree will form one contiguous block in
the sequence of visitations for the whole tree, starting and ending
with $v$. The Euler tour will not visit any part of $v$'s subtree
otherwise, so we simply split the Euler tour BST before the first
visit to $v$ and after the last visit to $v$. This gives us the
Euler tour tree of $v$'s subtree and two trees that cover the
sequence before and and after $v$'s subtree. We concatenate those
last two trees, and we're done.
\item[-] \textsc{{\textbf Link}}\textbf{(v,w)}\\
As a child of $w$, we need $v$'s subtree to be traversed immediately
after and immediately before visits to $w$. This can be
accomplished by splitting $w$'s Euler tour tree after the first
visit to $w$, and then concatenating the left piece, $v$'s Euler
tree, a new tree of a singleton node of just $w$, and the right
piece. (It is also equivalent to choose to do the first split
immediately before the last visit to $w$ and to insert the singleton
node on the opposite side of $v$'s Euler tour tree because children
are not ordered in the represented tree.)
\end{itemize}
The above operations only use the basic BST operations that take
$O(\lg n)$ using any common balanced BST data structure, so all of
them run in $O(\lg n)$ time.
\section{Dynamic Graph Problems}
We wish to maintain an undirected graph subject to vertex insertion
and deletion, and edge insertion and deletion. Deleting a vertex also
deletes all incident edges (or alternatively the vertex must have no
incident edges to be deleted). Since a data structure that is never
observed can perform any operation in zero time, we will also say that
the graph can be queried in some way.
\subsection{Connectivity}
The query type we focus on in this lecture is connectivity. Given a graph
$G$, we wish to be able to handle queries of the form \textsc{\textbf
Connected}\textbf{(v,w)}, which returns true if and only if there is a path
from $v$ to $w$ in $G$. Also, we would like to handle queries of the form \textsc{\textbf
Connected}\textbf{($G$)}, which returns whether or not $G$ is connected.
We have already seen how to achieve $O(\lg n)$ time per operation when $G$
is a tree --- use the link-cut tree data-stracture and compare \textsc{\textbf
Findroot}\textbf{(v)} with \textsc{\textbf
Findroot}\textbf{(w)}. There have been several other results for graphs
that are not forests:
\begin{itemize}
\item[-] $O(\lg n)$ update and query for plane graphs \cite{eppstein}.
\item[-] $O(\lg n(\lg\lg n)^3)$ update\\ $O(\lg n / \lg\lg\lg n)$ query
\cite{thorupstoc}
\item[-] $O(\lg^2 n)$ update\\$O(\lg n / \lg\lg n)$ query \cite{hlt}
\item[-] $O(\lg n\cdot x)$ update requires $\Omega(\lg n / \lg x)$ query for
$x > 1$ \cite{patdem}.
\item[-] \textbf{OPEN: } Is $o(\lg n)$ update and $O(\poly(\lg n))$
query achievable?
\item[-] \textbf{OPEN: } Is $O(\lg n)$ update and query for general graphs achievable?
\end{itemize}
For the first result, a {\em plane} graph is similar to a {\em planar}
graph, but the planar embedding is fixed (we are not allowed to change
the embedding as we add edges to the graph). Also, all the bounds
above are amortized. There has only been one significant worst-case
bound proven, done in \cite{egin}. They showed that $O(\sqrt{n})$
update and $O(1)$ query is achievable. It is an open problem as to
whether we can achieve $O(polylog(n))$ updates and queries in the worst
case. Note that the bounds in \cite{thorupstoc,hlt,egin} are in a
sense optimal in that they achieve the trade-off bounds shown in
\cite{patdem}.
There are some modifications we can play in the types of operations we
allow to be done to our graph. For {\em incremental} dynamic graphs, we
are never allowed to delete edges or vertices. For this problem, we know
that we can achieve $O(\alpha(n))$ update and query by using the union-find
data structure ($\alpha(n)$ being the inverse Ackermann function). It is
also known that achieving $\Theta(x)$ update requires $\Theta(\lg n / \lg
x)$ for $x > 1$. This has also been achieved. For {\em decremental}
dynamic graphs we allow deletions but no insertions. In \cite{thorupjacm}
it was shown how to achieve $O(m\lg n + n\poly(\log n))$ for all updates,
and $O(1)$ query (where $m$ is the number of edges).
\section{Dynamic Connectivity}
Now we will focus on the explanation of the dynamic connectivity
algorithm described in \cite{hlt}. The high level idea is that we
will store the spanning forest using Euler tour trees. We will
hierarchically divide the connected components into $\lg n$ levels of
spanning forests of subgraphs.
The {\em level} of an edge is some integer between $0$ and $\lg n$,
and it only decreases over time. $G_i$ is the subgraph consisting of
edges that are at level $i$ or less. Thus $G_{\lg n}=G$. The $F_i$
will also be the spanning forest of $G_i$. We will keep two
invariants during the execution of the algorithm:
\paragraph{Invariant 1: }
Every connected component of $G_i$ has at most $2^i$ vertices.
\paragraph{Invariant 2: } $F_0\subseteq F_1\subseteq\ldots\subseteq
F_{\lg n}$. In other words, $F_i = F_{\lg n}\bigcap G_i$, and $F_{\lg
n}$ is the minimum spanning forest of $G_{\lg n}$, using edge levels
as weights.
We will also keep an adjacency matrix for each $G_i$. There are three
operations we would like our data structure to support:
\begin{itemize}
\item[-] \textsc{{\textbf Insert}}\textbf{(e=(v,w))}\\
We will set the level of $e$ to $\lg n$ and update the adjacency
lists of $v$ and $w$. Also, if $v$ and $w$ are in separate connected
components of $F_{\lg n}$ then we add $e$ to $F_{\lg n}$ (we can tell
this by calling \textsc{Findroot} on $v$ and $w$.
\item[-] \textsc{{\textbf Delete}}\textbf{(e=(v,w))}\\
First we remove $e$ from the adjacency lists of $v$ and $w$. Then,
we do the following:
\begin{itemize}
\item[-] if $e$ is in $F_{\lg n}$
\begin{itemize}
\item[-] delete $e$ from $F_i$ for $i\geq$level$(e)$
\item[-] look for a replacement edge to reconnect $v$ and $w$
\begin{itemize}
\item[-] the replacement edge cannot be at a level less
than level$(e)$ by invariant 2 (each $F_i$ is a minimum
spanning forest).
\item[-] We will start searching for a replacement edge
at level$(e)$ to preserve the second invariant.
\end{itemize}
\item[-] for $i=$level$(e)$,...,$\lg n$:
\begin{itemize}
\item[-] Let $T_v$ be the tree containing $v$ and $T_w$
the tree containing $w$.
\item[-] Relabel $v$ and $w$ so that $|T_v|\leq |T_w|$.
\item[-] By invariant 1, we know that $|T_v| +
|T_w|\leq 2^i\Rightarrow |T_v|\leq 2^{i-1}$ (This
means we can afford to push all the edges of $T_v$
down to level $i-1$).
\item[-] for each edge $e'=(x,y)$ with $x$ in $T_v$ and
level$(e')=i$
\item[] \ \ \ \ - if $y$ is in $T_w$: add $(x,y)$ to
$F_i,F_{i+1},\ldots,F_{\lg n}$ and stop.
\item[] \ \ \ \ - else set level$(e')$ to $i-1$.
\end{itemize}
\end{itemize}
\end{itemize}
\item[-] \textsc{{\textbf Connected}}\textbf{(v,w)}
Instead of storing the nodes in $F_{\lg n}$ using a balanced binary
search tree, we can use a B-tree with a branching factor of $O(\lg
n)$, which gives us a query time of $O(\lg n / \lg\lg n)$. Using
this tree means that the update time is $O(\lg^2n / \lg\lg n)$ for
our other two operations.
\end{itemize}
We have to augment our Euler tour trees to do delete so that we can test if
$|T_v| \leq |T_w|$ in $O(1)$ time (this part is standard and easy). We
also would like to know for each node $v$ in the minimum spanning forest
$F_i$ whether or not $v$'s subtree contains any nodes incident to level-$i$
edges. We can find the next level-$i$ edge incident to a node in $T_v$ in
$O(\lg n)$ time using successor, jumping over empty subtrees. There is
also an $O(\lg n)$ charge to each edge's decrease in level.
The total cost of delete is then $O(\lg^2 n)$ since each inserted edge is
charged $O(\lg n)$ times. The remaining part of delete also takes $O(\lg^2
n)$ time since we have to delete the edge from at most $O(\lg n)$ levels,
and deleting from each level takes $O(\lg n)$.
\subsection{Other Dynamic Graph Problems}
\subsubsection{Minimum Spanning Forest}
Here we would like an MST for each connected component of the graph,
each represented as a dynamic tree. An amortized solution to this was
described in \cite{hlt} that achieves $O(\lg^4 n)$ update. A worst
case bound of $O(\sqrt{n})$ for general graphs was achieved in
\cite{egin}, and $O(\lg n)$ for plane graphs was achieved in
\cite{eitjwy}. The problem of determining graph bipartiteness can
also be solved by reduction to the minimum spanning forest problem.
\subsubsection{k-connectivity (vertex or edge)}
A pair of vertices $(v,w)$ is {\em k-connected} if there are $k$
vertex-disjoint (or edge-disjoint) paths from $v$ to $w$. A graph is
k-connected if each vertex pair in the graph is k-connected. To
determine whether the whole graph is k-connected can be solved as a
max-flow problem. An $O(\sqrt{n}\poly(\lg n))$ algorithm for
$O(\poly(\lg n))$-edge-connectivity was shown in \cite{thorup2001}.
There have been many results for k-connectivity between a single pair
of vertices:
\begin{itemize}
\item[-] $O(\poly(\lg n))$ for $k=2$ \cite{hlt}
\item[-] $O(\lg^2 n)$ for planar, decremental graphs \cite{italiano}
\end{itemize}
The above are amortized bounds. There have also been worst case bounds
shown in \cite{egin}:
\begin{itemize}
\item[-] $O(\sqrt{n})$ for 2-edge-connectivity
\item[-] $O(n)$ for 2-vertex-connectivity and 3-vertex-connectivity
\item[-] $O(n^{2/3})$ for 3-edge-connectivity
\item[-] $O(n\alpha(n))$ for $k=4$
\item[-] $O(n\lg n)$ for $O(1)$-edge-connectivity
\item[-] \textbf{OPEN: } Is $O(\poly(\lg n))$ achievable for $k=O(1)$?
Or perhaps $k=\poly(\lg n)$?
\end{itemize}
\subsubsection{Planarity Testing}
For this type of query, we would like to ask if inserting some edge
$e=(v,w)$ into our graph violates planarity. It was shown in
\cite{gis} that we can do this in $O(n^{2/3})$. In \cite{ipr} it is
shown how to achieve $O(\lg^2 n)$ worst case for a fixed embedding.
La Poutr\'{e} showed an $O(\alpha(m,n)\cdot m+n)$ algorithm for a
total of $m$ operations with an incremental graph in \cite{poutre}.
Planarity is just one type of minor-closed property, i.e. by
Kuratowski's theorem we know that a graph is planar if and only if it
does not contain a subgraph isomorphic to $K_5$ or to $K_{3,3}$. In
general, it is an open problem to test whether or not a graph has a
fixed minor. A minor of $G$ is a graph that can be obtained by a
sequence of edge deletions and contractions. A property is
minor-closed if $G$ having the property implies that any minor of $G$
must also have the property. It was shown in \cite{robertson} that
$G$ has a certain minor-closed property if and only if it excludes a
finite set of minors.
%\bibliography{mybib}
\bibliographystyle{alpha}
\begin{thebibliography}{77}
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Monika Rauch Henzinger, Valerie King: Randomized dynamic graph algorithms
with polylogarithmic time per operation. STOC 1995: 519-527
\bibitem{eppstein}
David Eppstein, Giuseppe F. Italiano, Roberto Tamassia, Robert Endre
Tarjan, Jeffery Westbrook, Moti Yung: Maintenance of a Minimum Spanning Forest in a Dynamic Plane Graph. J. Algorithms 13(1): 33-54 (1992)
\bibitem{thorupstoc}
Mikkel Thorup: Near-optimal fully-dynamic graph connectivity. STOC 2000: 343-350
\bibitem{hlt}
Jacob Holm, Kristian de Lichtenberg, Mikkel Thorup: Poly-logarithmic
deterministic fully-dynamic algorithms for connectivity, minimum spanning
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\bibitem{patdem}
Mihai Patrascu, Erik D. Demaine: Lower bounds for dynamic
connectivity. STOC 2004: 546-553
\bibitem{egin}
David Eppstein, Zvi Galil, Giuseppe F. Italiano, Amnon Nissenzweig:
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\bibitem{thorupjacm}
Mikkel Thorup: Decremental Dynamic Connectivity. J. Algorithms 33(2):
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\bibitem{eitjwy}
David Eppstein, Giuseppe F. Italiano, Roberto Tamassia, Robert Endre
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\bibitem{thorup2001}
Mikkel Thorup: Fully-dynamic min-cut. STOC 2001: 224-230
\bibitem{italiano}
Dora Giammarresi, Giuseppe F. Italiano: Decremental 2- and 3-Connectivity
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\bibitem{gis}
Zvi Galil, Giuseppe F. Italiano, Neil Sarnak: Fully Dynamic Planarity
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Giuseppe F. Italiano, Johannes A. La Poutr\'{e}, Monika Rauch: Fully Dynamic
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Johannes A. La Poutr\'{e}: Alpha-algorithms for incremental planarity
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Neil Robertson, Paul D. Seymour: Graph Minors. XX. Wagner's conjecture. J. Comb. Theory, Ser. B 92(2): 325-357 (2004)
\end{thebibliography}
\end{document}