# Reasoning About Algorithms

# COMP 382

## BSTs §

Assumption: keys are linearly ordered.

Let $x$ and $y$ be nodes in the tree.

- If node $y$ is in the left subtree of $x$, then $y_{key} \leq x_{key}$
- If node $y$ is in the right subtree of $x$, then $y_{key} \geq x_{key}$

Every subtree of a BST is a BST.

### Traversal §

```
in_order(x):
if x != null:
in_order(x.left)
print(x.key)
in_order(x.right)
```

- preorder: root, left, right
- inorder: left, root, right
- postorder: left, right, root

Running time: $\Theta(n)$ where $n$ is the tree size

### Search §

```
tree_search(x, k):
if x == null or k == x.key
return x
if k < x.key:
return tree_search(x.left, k)
return tree_serach(x.right, k)
```

$O(h)$, where $h$ is the height of the tree. $h$ is

```
tree_search(x, k):
while x != null and k != x.key:
if k < x.key:
x = x.left
else:
x = x.right
return x
```

### Successor §

Pre-condition: `x != null`

```
tree_successor(x):
if x.right != null:
return tree_min(x)
y = x.p
while y != null and x == y.right:
x = y
y = y.p
return y
```

## Red-Black Trees §

- every node is either red or black
- the root is black
- every leaf (NIL) is black
- the children of a red node must be black
- for every node $u$, all simple paths from $u$ to descendent leaves have the same number of black nodes.

properties 4 and 5 together ensure that the tree is “roughly balanced”.

## AVL Trees §

Binary search trees that are balanced.

For every node $x$ with a left subtree $l$ and a right subtree $r$, $|\textrm{height}(r) - \textrm{height}(l)| \leq 1$. The heights of subtrees $l$ and $r$ differ by at most 1.

Balance factor $\textrm{balance}(x) = \textrm{height}(x.r) - \textrm{height}(x.l)$

## B-Trees §

Assumption: each node is stored on a separate disk page.

The branching factor of B trees can be large.

A node with $n$ keys has $n + 1$ children.

A rooted tree:

- every node
`x`

has attributes- number of keys stored
`x.n`

- the keys
`x.key[1] <= x.key[2] <= ... <= x.key[x.n]`

(ordered) - boolean value
`x.leaf`

- number of keys stored
- every internal node
`x`

has`x.n + 1`

pointers`x.c[1]`

,`x.c[2]`

, …,`x.c[x.n + 1]`

to its children - the keys separate the ranges of keys for subtrees
- all leaves have the same depth, which is equal to the height
`h`

of the tree - a b-tree has a parameter $t \geq 2$ called the minimum degree
- each node (except root) has at least $t - 1$ keys. each internal node (except root) has at least $t$ children.
- each node has at msot $2t - 1$ keys

Consider a non-empty B-tree $T$ with minimum degree $t$. The height of $T$ is $h$.

```
# assumption: x.node has already been read from disk k.key
b_tree_search(x, k):
i = 1
while i <= x.n and k > x.key[i]:
i++
if i <= x.n and k == x.key[i]:
return (x, i)
else if x.leaf:
return null
else:
disk_read(x, c[i])
return b_tree_search(x.c[i], k)
```

$O(h)$ reads from disk and $O(t \cdot h)$ running time.

## Hash Tables §

Data structure that stores keys with additional satellite data

### The universe is small §

Direct address table: array `T[0..m-1]`

.

- If key
`k`

is not present,`T[k] = NIL`

- If key
`k`

is present and has value`val`

,`T[k] = val`

### The universe is large; the set $k$ of present keys is small §

Memory: $\Theta(|k|)$

Search: $O(1)$ time in the average case

Components:

- table
`T[0..m-1]`

where typically m \llt |U| - hash function $h : U \rightarrow \{0, 1, \cdots, m - 1\}$
- key $k$ hashes to the slot $h(k)$

values stored in `T[h(k)]`

. collisions unavoidable since $m < |U|$.

chaining: elements that hash to the same value are placed into a linked list.

`insert(T, x)`

: insert`x`

at the head of list`T[h(x.key)]`

`search(T, k)`

: search for`k`

in the list`T[h(k)]`

`delete(T, x)`

: delete`x`

from the list`T[h(x.key)]`

worst case for search: all $n$ keys hash to the same slot. searching will be $\Theta(n)$.

[Assumption: simple, uniform hashing]

We insert keys $k_1, k_2, \cdots, k_n$ into a hash table with $m$ slots

$P(h(i) = j) = \frac{1}{m}$

$n_j$ length of the list `T[j]`

load factor

## Greedy Algorithms §

### Interval Partitioning §

The *depth* of a set of intervals is the maximum number of intervals that contain any given time.

Number of classrooms needed ≥ Depth

- Consider intervals in increasing order of start time.
- Assign any compatible classroom.

Proof strategy: Find a structural bound asserting that every possible solution must have been at least a certain value, then show that the algorithm always achieves this bound.

### Scheduling to Minimize Lateness §

- Satisfy all requestts
- Allow certain requests to run late
- Minimize the max lateness

Formally:

- Single resource processing one job at a time
- Job $j$ requires $t_j$ time and is due at $d_j$
- If $j$ starts at $s_j$, it finishes at $f_j = s_j + t_j$
- Lateness: $l_j = \max(0, f_j - d_j)$
- Schedule to minimize max lateness $L = \max l_j$

“Earliest deadline first” is optimal.

greedy-is-optimal proof techniques:

- greedy always stays ahead
- structural
- exchange argument

## Minimum Spanning Tree §

**Spanning tree**: Given a connected, undirected graph $G = (V, E)$ with real edge weights $c_e$, a spanning tree $T$ contains a subset of $E$ that connects all vertices and whose total weight $w(T) = \Sigma_{u, v \in T} w(u, v)$. $T$ is undirected (no root) and is, by definition, acyclic.

When $G$ has $n$ nodes, $T$ has $n - 1$ edges

**Minimum spanning tree** (MST): a spanning tree of $G$ with lowest total weight $w$

Properties of MSTs:

- every connected, undirected graph $G$ has at least one MST
- having distinct edge costs is a sufficient (but not necessary) condition for $G$ to have a unique MST
- between every pair of nodes in a spanning tree, there is exactly one path

A **cut** $(S, V - S)$ or $S$ of $G$ is a partition of the nodes $V$ into the two non-empty sets $S$ and $V - S$

**Cut property**: For a cut $S$ in $G$, the MST contains the min-cost crossing edge $e$

**Cycle property**: For a cycle $C$ in $G$ the MST does not contain the max-cost edge in $C$

### Kruskal’s algorithm §

- $T = \varnothing$
- Sort the edges in ascending order of cost
- Iterate over the edges in order
- if adding edge $e$ to $T$ creates a cycle, discard $e$ (cycle property)
- Else, insert $e = (u, v)$ into $T$

We need to be able to efficiently check if adding an edge will form a cycle. We need to maintain a set for each connected component.

As an edge $e = (u, v)$ is considered:

- find the CCs containing $u$ and $v$
- if they’re different there’s no path between $u$ and $v$
- add edge $e$
- merge the components

- if they’re the same there’s already a path between $u$ and $v$. don’t add.

Union-Find.

`make_set`

: construct a singleton set`find(u)`

: find the set containing`u`

`union(a, b)`

: merge the two subsets`a`

and`b`

into a single set

Best implemented with disjoint set forests.

- union by rank
- path compression

Kruskal takes $O(m \log m)$ for sorting and $O(m \alpha(m, n)$ for find operations. $\alpha(m, n)$ grows very slowly (effectively a small constant).

## Dynamic Programming §

### Edit Distance §

- $Dist(\epsilon, \epsilon) = 0$
- $Dist(X, \epsilon) = |X|$
- $Dist(\epsilon, Y) = |Y|$

$Dist(Xx, Yy) = \min\begin{pmatrix}Dist(X, Y) + cost\binom{x}{y} \\ Dist(Xx, Y) + 1 \\ Dist(X, Yy) + 1\end{pmatrix}$

### LCS §

- $LCSLength(\epsilon, \epsilon) = 0$
- $LCSLength(X, \epsilon) = 0$
- $LCSLength(Xz, Yz) = LCSLength(X, Y) + 1$

Finally, for $x \neq y$:

$LCSLength(Xx, Yy) = \max\begin{pmatrix}LCSLength(Xx, Y) \\ LCSLength(X, Yy)\end{pmatrix}$

### Optimal BST §

$\sum p_i + \sum q_i = 1$

$E[\textrm{search cost}] = \sum p_i \cdot ({depth}_T(k_i) + 1) + \sum q_i \cdot ({depth}_T(d_i) + 1)$

## Randomized Algorithms §

### Selection §

Find the $k$-th largest element in a given array.

$T(n) = (n - 1) + \frac{1}{n} \sum_{i = 1}^{n - 1} T(i)$

### Quicksort §

```
def quicksort(S):
if |S| <= 3:
sort S
return S
choose a splitter x uniformly at random
for each element S_i of S:
if S_i <= x: put S_i in S_
if S_i > x: put S_i in S__
quicksort(S_)
quicksort(S__)
return S_ + S__
```

$T(n) = (n - 1) + \frac{1}{n} \sum_{i = 1}^n [T(i - 1) + T(n - i)]$

## Maximum Flow §

A flow graph $G = (V, E)$ is a directed, weighted graph in which each edge has a capacity $c(e) \geq 0$. source $s$, sink $t$.

- every vertex lies on some path from $s$ to $t$
- $(u, v) \in E \implies (v, u) \not\in E$
- no parallel edges

A flow is a real-valued function $f: V \times V \rightarrow \mathbb{R}$:

- [capacity] for each edge $e$, $0 \leq f(e) \leq c(e)$
- [conservation] for each $v \in V - \{s, t\}$, $\sum_{e \textrm{ into }v} f(e) = \sum_{e \textrm{ out of }v} f(e)$

The *value of flow* $f$ is $\sum_{e \textrm{ out of }s} f(e)$

**Minimum cut**: partition which minimizes the capacity across the two parts

Max flow = Min cut

For a flow $f$ and an $s$-$t$ cut $(A, B)$, the *net-flow across the cut* is equal to the net amount leaving $A$.

$\sum_{e \textrm{ out of }A} f(e) = \sum_{e \textrm{ into }A} f(e)$

**Flow value lemma**: the net flow across the cut is equal to the value of the flow.

$\sum_{e \textrm{ out of }A} f(e) = \sum_{e \textrm{ into }A} f(e) = v(f)$

The flow is bounded by cut capacity. For a flow $f$ and a cut $(A, B)$, the value of the flow is at most the capacity of the cut.

$v(f) \leq cap(A, B)$