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# 26.4 Push-relabel algorithms

## 26.4-1

Prove that, after the procedure $\text{INITIALIZE-PREFLOW}(G, S)$ terminates, we have $s.e \le -|f^*|$, where $f^*$ is a maximum flow for $G$.

We apply the definition of excess flow (equation $\text{(26.14)}$) to the initial preflow $f$ created by $\text{INITIALIZE-PREFLOW}$ (equation $\text{(26.15)}$) to obtain

\begin{aligned} e(s) & = \sum_{v \in V} f(v, s) - \sum_{v \in V} f(s, v) \\ & = 0 - \sum_{v \in V} c(s, v) \\ & = -\sum_{v \in V} c(s, v). \end{aligned}

Now,

\begin{aligned} -|f^*| & = \sum_{v \in V} f^*(v, s) - \sum_{v \in V} f^*(s, v) \\ & \ge 0 - \sum_{v \in V} c(s, v) \qquad \text{(since f^*(v, s) \ge 0 and f^*(s, v) \le c(s, v))} \\ & = e(s). \end{aligned}

## 26.4-2

Show how to implement the generic push-relabel algorithm using $O(V)$ time per relabel operation, $O(1)$ time per push, and $O(1)$ time to select an applicable operation, for a total time of $O(V^2E)$.

We must select an appropriate data structure to store all the information which will allow us to select a valid operation in constant time. To do this, we will need to maintain a list of overflowing vertices. By Lemma 26.14, a push or a relabel operation always applies to an overflowing vertex. To determine which operation to perform, we need to determine whether $u.h = v.h + 1$ for some $v \in N(u)$. We'll do this by maintaining a list $u.high$ of all neighbors of $u$ in $G_f$ which have height greater than or equal to $u$. We'll update these attributes in the $\text{PUSH}$ and $\text{RELABEL}$ functions. It is clear from the pseudocode given for $\text{PUSH}$ that we can execute it in constant time, provided we have maintain the attributes $\delta_f(u, v)$, $u.e$, $c_f(u, v)$, $(u, v).f$ and $u.h$. Each time we call $\text{PUSH}(u, v)$ the result is that $u$ is no longer overflowing, so we must remove it from the list.

Maintain a pointer $u.overflow$ to $u$'s position in the overflow list. If a vertex $u$ is not overflowing, set $u.overflow = \text{NIL}$. Next, check if $v$ became overflowing. If so, set $v.overflow$ equal to the head of the overflow list. Since we can update the pointer in constant time and delete from a linked list given a pointer to the element to be deleted in constant time, we can maintain the list in $O(1)$.

The $\text{RELABEL}$ operation takes $O(V)$ because we need to compute the minimum $v.h$ from among all $(u, v) \in E_f$, and there could be $|V| - 1$ many such $v$. We will also need to update $u.high$ during $\text{RELABEL}$. When $\text{RELABEL}(u)$ is called, set $u.high$ equal to the empty list and for each vertex $v$ which is adjacent to $u$, if $v.h = u.h + 1$, add $u$ to the list $v.high$. Since this takes constant time per adjacent vertex we can maintain the attribute in $O(V)$ per call to relabel.

## 26.4-3

Prove that the generic push-relabel algorithm spends a total of only $O(VE)$ time in performing all the $O(V^2)$ relabel operations.

Each time we call $\text{RELABEL}(u)$, we examine all edges $(u, v) \in E_f$. Since the number of relabel operations is at most $2|V| - 1$ per vertex, edge $(u, v)$ will be examined during relabel operations at most $4|V| - 2 = O(V)$ times (at most $2|V| - 1$ times during calls to $\text{RELABEL}(u)$ and at most $2|V| - 1$ times during calls to $\text{RELABEL}(v)$). Summing up over all the possible residual edges, of which there are at most $2|E| = O(E)$, we see that the total time spent relabeling vertices is $O(VE)$.

## 26.4-4

Suppose that we have found a maximum flow in a flow network $G = (V, E)$ using a push-relabel algorithm. Give a fast algorithm to find a minimum cut in $G$.

We can find a minimum cut, given a maximum flow found in $G = (V, E)$ by a push-relabel algorithm, in $O(V)$ time. First, find a height $\hat h$ such that $0 < \hat h < |V|$ and there is no vertex whose height equals $\hat h$ at termination of the algorithm. We need consider only $|V| - 2$ vertices, since $s.h = |V|$ and $t.h = 0$. Because $\hat h$ can be one of at most $|V| - 1$ possible values, we know that for at least one number in $1, 2, \ldots, |V| - 1$, there will be no vertex of that height. Hence, $\hat h$ is well defined, and it is easy to find in $O(V)$ time by using a simple boolean array indexed by heights $1, 2, \ldots, |V| - 1$.

Let $S = \{u \in V: u.h > \hat h\}$ and $T = \{v \in V: v.h < \hat h\}$. Because we know that $s.h = |V| > \hat h$, we have $s \in S$, and because $t.h = 0 < \hat h$, y we have $t \in T$, as required for a cut.

We need to show that $f(u, v) = c(u, v)$, i.e., that $(u, v) \notin E_f$, for all $u \in S$ and $v \in T$. Once we do that, we have that $f(S, T) = c(S, T)$, and by Corollary 26.5, $(S, T)$ is a minimum cut.

Suppose for the purpose of contradiction that there exist vertices $u \in S$ and $v \in T$ such that $(u, v) \in E_f$. Because $h$ is always maintained as a height function (Lemma 26.16), we have that $u.h \le v.h + 1$. But we also have $v.h < \hat h < u.h$, and because all values are integer, $v.h \le u.h - 2$. Thus, we have $u.h \le v.h + 1 \le u.h - 2 + 1 = u.h - 1$, which gives the contradiction that $u.height \le u.height - 1$. Thus, $(S, T)$ is a minimum cut.

## 26.4-5

Give an efficient push-relabel algorithm to find a maximum matching in a bipartite graph. Analyze your algorithm.

First, construct the flow network for the bipartite graph as in the previous section. Then, we relabel everything in $L$. Then, we push from every vertex in $L$ to a vertex in $R$, so long as it is possible.

Keeping track of those that vertices of $L$ that are still overflowing can be done by a simple bit vector. Then, we relabel everything in R and push to the last vertex. Once these operations have been done, The only possible valid operations are to relabel the vertices of $L$ that weren't able to find an edge that they could push their flow along, so could possibly have to get a push back from $R$ to $L$. This continues until there are no more operations to do. This takes time of $O(V(E + V))$.

## 26.4-6

Suppose that all edge capacities in a flow network $G = (V, E)$ are in the set $\{1, 2, \ldots, k\}$. Analyze the running time of the generic push-relabel algorithm in terms of $|V|$, $|E|$, and $k$. ($\textit{Hint:}$ How many times can each edge support a nonsaturating push before it becomes saturated?)

The number of relabel operations and saturating pushes is the same as before. An edge can handle at most $k$ nonsaturating pushes before it becomes saturated, so the number of nonsaturating pushes is at most $2k|V||E|$. Thus, the total number of basic operations is at most $2|V|^2 + 2|V||E| + 2k|V||E| = O(kVE)$.

## 26.4-7

Show that we could change line 6 of $\text{INITIALIZE-PREFLOW}$ to

 1  6 s.h = |G.V| - 2 

without affecting the correctness or asymptotic performance of the generic pushrelabel algorithm.

If we set $s.h = |V| - 2$, we have to change our definition of a height function to allow $s.h = |V| - 2$, rather than $s.h = |V|$. The only change we need to make to the proof of correctness is to update the proof of Lemma 26.17. The original proof derives the contradiction that $s.h \le k < |V|$, which is at odds with $s.h = |V|$. When $s.h = |V| - 2$, there is no contradiction.

As in the original proof, let us suppose that we have a simple augmenting path $\langle v_0, v_1, \ldots, v_k \rangle$, where $v_0 = s$ and $v_k = t$, so that $k < |V|$. How could $(s, v_1)$ be a residual edge? It had been saturated in $\text{INITIALIZE-PREFLOW}$, which means that we had to have pushed some flow from $v_1$ to $s$. In order for that to have happened, we must have had $v_1.h = s.h + 1$. If we set $s.h = |V| - 2$, then $v_1.h$ was $|V| - 1$ at the time. Since then, $v_1.h$ did not decrease, and so we have $v_1.h \ge |V| - 1$. Working backwards over our augmenting path, we have $v_{k - i}.h \le t.h + i$ for $i = 0, 1, \ldots, k$. As before, because the augmenting path is simple, $k < |V|$. Letting $i = k - 1$, we have $v_1.h \le t.h + k - 1 < 0 + |V| - 1$. We now have the contradiction that $v_1.h \ge |V| - 1$ and $v_1.h < |V| - 1$, which shows that Lemma 26.17 still holds.

Nothing in the analysis changes asymptotically.

## 26.4-8

Let $\delta_f(u, v)$ be the distance (number of edges) from $u$ to $v$ in the residual network $G_f$. Show that the $\text{GENERIC-PUSH-RELABEL}$ procedure maintains the properties that $u.h < |V|$ implies $u.h \le \delta_f(u, t)$ and that $u.h \ge |V|$ implies $u.h - |V| \le \delta_f(u, s)$.

We'll prove the claim by induction on the number of push and relabel operations. Initially, we have $u.h = |V|$ if $u = s$ and $0$ otherwise. We have $s.h - |V| = 0 \le \delta_f(s, s) = 0$ and $u.h = 0 \le \delta_f(u, t)$ for all $u \ne s$, so the claim holds prior to the first iteration of the while loop on line 2 of the $\text{GENERIC-PUSH-RELABEL}$ algorithm.

Suppose that the properties have been maintained thus far. If the next iteration is a nonsaturating push then the properties are maintained because the heights and existence of edges in the residual network are preserved. If it is a saturating push then edge $(u, v)$ is removed from the residual network, which increases both $\delta_f(u, t)$ and $\delta_f(u, s)$, so the properties are maintained regardless of the height of $u$.

Now suppose that the next iteration causes a relabel of vertex $u$. For all $v$ such that $(u, v) \in E_f$ we must have $u.h \le v.h$. Let $v' = \min\{v.h \mid (u,v) \in E_f\}$. There are two cases to consider.

• First, suppose that $v.h < |V|$. Then after relabeling we have

$$u.h = 1 + v'.h \le 1 + \min_{(u, v)} \in E_f \delta_f(v, t) = \delta_f(u, t).$$

• Second, suppose that $v'.h \ge |V|$. Then after relabeling we have

$$u.h = 1 + v'.h \le 1 + |V| + \min_{(u, v)} \in E_f \delta_f(v, s) = \delta_f(u, s) + |V|,$$

which implies that $u.h - |V| \le \delta_f(u, s)$.

Therefore, the $\text{GENERIC-PUSH-RELABEL}$ procedure maintains the desired properties.

## 26.4-9 $\star$

As in the previous exercise, let $\delta_f(u, v)$ be the distance from $u$ to $v$ in the residual network $G_f$. Show how to modify the generic push-relabel algorithm to maintain the property that $u.h < |V|$ implies $u.h = \delta_f(u, t)$ and that $u.h \ge |V|$ implies $u.h - |V| = \delta_f(u, s)$. The total time that your implementation dedicates to maintaining this property should be $O(VE)$.

What we should do is to, for successive backwards neighborhoods of $t$, relabel everything in that neighborhood. This will only take at most $O(VE)$ time (see 26.4-3). This also has the upshot of making it so that once we are done with it, every vertex's height is equal to the quantity $\delta_f(u, t)$. Then, since we begin with equality, after doing this, the inductive step we had in the solution to the previous exercise shows that this equality is preserved.

## 26.4-10

Show that the number of nonsaturating pushes executed by the $\text{GENERIC-PUSH-RELABEL}$ procedure on a flow network $G = (V, E)$ is at most $4|V|^2|E|$ for $|V| \ge 4$.

Each vertex has maximum height $2|V| - 1$. Since heights don't decrease, and there are $|V| - 2$ vertices which can be overflowing, the maximum contribution of relabels to $\Phi$ over all vertices is $(2|V| - 1)(|V| - 2)$. A saturating push from $u$ to $v$ increases $\Phi$ by at most $v.h \le 2|V| - 1$, and there are at most $2|V||E|$ saturating pushes, so the total contribution over all saturating pushes to $\Phi$ is at most $(2|V| - 1)(2|V||E|)$. Since each nonsaturating push decrements $\Phi$ by at least on and $\Phi$ must equal zero upon termination, we must have that the number of nonsaturating pushes is at most

$$(2|V| - 1)(|V| - 2) + (2|V| - 1)(2|V||E|) = 4|V|^2|E| + 2|V|^2 - 5|V| + 3 - 2|V||E|.$$

Using the fact that $|E| \ge |V| - 1$ and $|V| \ge 4$ we can bound the number of saturating pushes by $4|V|^2|E|$.