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The 2-core is what remains of a finite simple graph after repeatedly removing every vertex with fewer than two surviving graph neighbours. It strips away dangling trees while keeping the part that can support itself under this local rule. Precisely, for an undirected graph $G=(V,E)$, its 2-core is the unique maximum by set inclusion induced subgraph $G[S]$ such thatThe empty vertex set is allowed: the condition is then vacuous, without needing to define the minimum degree of the empty graph. "Maximum by inclusion" means that $S$ contains every other vertex set satisfying the condition. It is stronger than merely saying that one chosen set cannot be enlarged: the 2-core simultaneously contains all possible qualifying induced subgraphs. In this particular family, every inclusion-maximal qualifying set does equal that unique maximum; the union argument in why the 2-core is unique explains why. ^[ref]

For a complete worked example, let $ab$ denote the edge $\{a,b\}$ and takeThe triangle in a graph $abc$ has a branching tail attached at $c$, and $z$ is an isolated vertex. In alphabetical order the initial degrees of vertices are $(2,2,3,3,2,1,1,0)$. Deleting $f$ turns $e$ into a leaf of a graph; deleting $g$ leaves $d$ with degree of a vertex two; deleting $z$ changes no other degree of a vertex. Now delete $e$, which makes $d$ a leaf of a graph, and finally delete $d$. Exactly the three edges of $abc$ remain: $C_2(G)=G[\{a,b,c\}]$.

A forest has empty 2-core: every nonempty finite forest has an isolated vertex or a leaf of a graph, so deletion cannot stop at a nonempty remainder. In particular, an isolated edge, an isolated vertex, a path graph, and the empty graph all have empty 2-core. Every cycle graph is its own 2-core. Two disjoint triangles in a graph also form their own 2-core, so a 2-core need not be a connected graph; attaching a tree to either triangle changes neither surviving triangle.

The Aldous-Broder algorithm samples a uniform spanning tree of a finite connected graph by running a simple random walk from any initial vertex. Whenever the random walk first visits a new vertex, retain the edge by which it arrived. Stop when every vertex has been visited.

Each retained edge joins a previously unseen vertex to the already visited connected subgraph. It cannot create a graph cycle; at termination the retained edges form a spanning tree. The stopping time is the cover time, finite almost surely. The nontrivial conclusion is that every spanning tree has exactly the same probability, independently of the starting vertex.^[ref]

Repeated visits add no edges, which can make the method slow when only a few unvisited vertices remain. Wilson's algorithm instead attaches loop-erased random walks. On a recurrent infinite connected graph, the first-entrance construction continues indefinitely and gives the corresponding infinite-volume uniform spanning tree.

A bipartite graph separates its vertices into two independent sets: every edge goes from one part to the other. Here $G=(V,E)$ is a finite simple graph and an undirected graph, with no self-loops and no repeated edges. The two parts $X,Y\subseteq V$ satisfyThus neither part contains an edge. The parts need not have equal sizes, and some pairs across them may have no edge; requiring every cross pair gives a complete bipartite graph. The independent set and cross-edge definitions are equivalent.^[ref]

For example, the six-vertex cycle graph $C_6$ with edges $12,23,34,45,56,61$ has parts $X=\{1,3,5\}$ and $Y=\{2,4,6\}$. Drawing the same graph as a hexagon or in two columns changes neither its edges nor its bipartite property. An extra isolated vertex $z$ can go in either part.

We allow either part to be empty. Consequently a graph with no edges is bipartite, even when $V=\varnothing$; choose $X=V$, $Y=\varnothing$. Here "two colours" means two colours are available, with no requirement to use both. This convention explicitly includes the empty case, whereas definitions requiring two nonempty classes need an exception. For several connected components of a graph, choose the two parts separately in each graph component and take their unions. In a nonempty connected graph that is bipartite, fixing one vertex's colour forces all others along graph paths, so its two-colour assignment is unique up to swapping the colours. Each graph component, including an isolated vertex, can be swapped independently.

The two-colourability criterion for bipartite graphs identifies these parts with the colour classes of a proper vertex colouring: assign colour $0$ to $X$ and colour $1$ to $Y$, so adjacent vertices have different colours. More surprisingly, the odd-cycle characterization of bipartite graphs turns a global requirement into one precise obstruction:This statement is for the finite undirected graph and simple graph model above; a graph cycle has distinct vertices except for the repeated starting vertex. Along a graph cycle in a bipartite graph, colours alternate, so returning to the starting colour takes an even number of edges. Therefore any odd cycle, whether or not other edges join its vertices, prevents a bipartite split.^[ref]

Breadth-first search supplies the converse constructively. Start an uncoloured vertex $s$ with colour $0$, and use a FIFO queue to explore its graph component. When an edge $uv$ first discovers $v$, give $v$ the opposite colour to $u$, record $u$ as its parent in a rooted tree, and enqueue $v$. If both endpoints of any scanned edge already have the same colour, stop with a conflict. Restart from an uncoloured vertex until every graph component has been examined. If no conflict occurs, all edges cross the two colour classes, which are the required independent sets.^[ref]

A conflict also explains the failure. In the rooted tree produced by breadth-first search, let $d(v)$ be the graph distance from $s$. The colour of $v$ is its distance's integer parity. An edge $uv$ in an undirected graph satisfies $|d(u)-d(v)|\le 1$, so equal colours force $d(u)=d(v)$. Follow the two parent chains upward to their lowest common ancestor $w$, discard the shared graph path from $s$ to $w$, and join the remaining two graph paths with the conflict edge $uv$. The result is a genuine graph cycle withedges, an odd integer. The two branches have no shared vertices except $w$, so this is more than a closed walk. This odd cycle is a certificate that anyone can check directly, using only its listed vertices and edges.

In the example, breadth-first search from $s$ gives $d(w)=1$, $d(a)=d(b)=2$, and $d(u)=d(v)=3$. The conflict edge $uv$ closes the odd cycle $w,a,u,v,b,w$, with five edges. The common edge $sw$ is excluded. Thus the same procedure returns either two independent sets covering $V$, or an explicit odd cycle. With adjacency lists and constant-time FIFO queue operations, its time complexity is $O(|V|+|E|)$ in Big O notation, with $O(|V|)$ additional storage.^[ref]

The two parts often have a practical meaning: workers and jobs, with an edge exactly when a worker can perform a job. A matching in a graph chooses compatible worker–job pairs with no reused endpoint; merely being bipartite does not guarantee enough compatible pairs. Hall's marriage theorem characterizes when every worker can be matched, and the assignment problem adds costs to seek an optimal assignment. These connections make the bipartite structure useful after recognition, as the starting point for selecting edges under constraints.

If a finite simple graph is a connected graph, a bipartite graph, and a planar graph, with $n\ge 3$ vertices, then $m\le 2n-4$. Every facial boundary walk has even length and length at least three, hence at least four. Thus $4f\le 2m$; substitute the Euler formula for a connected planar graph.^[ref]

The square cycle graph attains equality: $4=2\cdot 4-4$. The three-vertex graph path also attains equality: $2=2\cdot 3-4$. Its repeated bridges in a graph illustrate why counting facial boundary walk length correctly matters. Neither bound applies unchanged to a multigraph: two vertices may support arbitrarily many parallel edges in a plane graph.

In log abilities, the negative Hessian matrix is a weighted graph Laplacian of the comparison graph. If that comparison graph is a connected graph and the abilities are finite, the quadratic form vanishes only for constant vectors. Thus the log-likelihood is strictly concave after fixing the common additive constant. If the directed graph of observed wins is a strongly connected directed graph, letting contrasts diverge forces at least one observed-win probability to zero, so the log-likelihood tends to negative infinity. A finite maximizer exists and is unique up to common scaling of abilities.

A branch set of a graph minor is the connected region representing one target vertex. For finite simple graphs $G$ and $H$, a witness to $H\preccurlyeq G$ assigns a nonempty set $B_x\subseteq V(G)$ to each $x\in V(H)$, such that the sets are pairwise disjoint, each induced subgraph $G[B_x]$ is a connected graph, and every $xy\in E(H)$ has a witnessing edge of $G$ between $B_x$ and $B_y$. Extra edges between branch sets of a graph minor are allowed: they can be deleted. Unused vertices are also allowed.^[ref]

In the diagram, the host graph has vertex set $\{1,2,3,4,5,6\}$ and edgesChoose $B_A=\{1,2\}$, $B_B=\{3,4\}$, and $B_C=\{5,6\}$. Their internal edges $12,34,56$ establish connectedness. The edges $23,45,61$ witness $AB,BC,CA$, respectively, so they form a model of the complete graph $K_3$. The additional edge $24$ also joins $B_A$ to $B_B$; it supplies another witness for the same target edge, not a fourth target edge. Contracting $12,34,56$ and merging duplicate edges gives exactly $\{AB,BC,CA\}$.

The same branch sets of a graph minor also witness the path graph $A,B,C$: retain the witnesses $23,45$, and delete $61$ to remove the unwanted $AC$ adjacency. This demonstrates why the definition requires an edge for every target adjacency, without forbidding other adjacencies in the host graph. For a complete graph $K_t$, there are no unwanted pairs: its model consists of $t$ disjoint connected branch sets of a graph minor with an edge between every pair.

To turn a model into an operation sequence, keep a spanning tree inside each $G[B_x]$, keep one witnessing edge for each edge of $H$, and delete all other edges and unused vertices. Contract each chosen spanning tree to one vertex. Disjointness prevents two target vertices from merging; connectedness provides the edge contractions that merge each individual region. Conversely, track the original vertices merged into every surviving vertex during a deletion-and-edge contraction sequence. They form disjoint connected branch sets of a graph minor, and each surviving edge comes from an original edge between its two regions. A list of these regions, their connecting trees, and their witnessing edges therefore gives a directly checkable certificate.

A bridge is an edge whose deletion increases the number of components of a graph. Equivalently, it belongs to no cycle; within its original component, deleting it makes the connected graph disconnected.

A 2-core can contain bridges in a graph, and even vertices that lie on no cycle in a graph. Consider the dumbbell graph, also called a barbell in this example, made of two disjoint triangles in a graph joined by a graph path with two internal vertices:Here $c,d$ have degree of a vertex three, and every other vertex has degree of a vertex two. No deletion is possible, so the whole graph is its 2-core. Yet $cx,xy,yd$ are bridges in a graph: deleting any one separates the two triangles. None belongs to a cycle in a graph, and $x,y$ belong to no cycle in a graph either. Keeping only edges on graph cycles would incorrectly discard the entire joining graph path.

Every cycle in a graph does survive pruning, because each of its vertices keeps its two cycle graph neighbours. The barbell shows that the converse fails: the 2-core can be larger than the union of all graph cycles. Its condition is local minimum degree, not protection against a single connection failing.

A biconnected component is a different object: a maximal connected subgraph with no cut vertex of its own. Under the block convention used here, the barbell has five biconnected components: its two triangles and its three individual joining edges, each with its endpoints. Its cut vertices are $c,x,y,d$. A 2-connected component with at least three vertices must remain connected after any single vertex is deleted; the whole barbell fails that test. An informal phrase such as "biconnected core" therefore needs its own definition and must not be substituted for 2-core. Even an isolated edge is a block under this convention, although its 2-core is empty.

The usefulness of the 2-core is this precise separation of local persistence from global robustness. In a connected graph with nonempty 2-core, the deleted vertices form a forest, with each of its connected components of a graph attached by exactly one edge to the 2-core. Otherwise a deleted graph cycle, or a graph path joining two surviving attachment points, would itself have survived. Stripping these attached trees exposes a smaller structure for studying graph cycles, graph excess and the graph kernel, while preserving the joining bridges in a graph that organize that structure.

A Cayley graph turns multiplication in a group into travel: the chosen generators of a group are the allowed one-step moves, and a word over an alphabet becomes a route. This lets us study a finitely generated group using graph distances, closed balls, and geodesics, rather than only equations between products.

Let $G$ be a group with a finite generating set of a group $S$ satisfying $S=S^{-1}$ and $1\notin S$. Such a symmetric subset of a group includes the inverse element of every allowed move. The Cayley graph ${\mathrm{Cay}}_S(G)$ has vertex set $G$ and joins $g$ to $gs$ for each $s\in S$. We choose right multiplication for these moves. Reversing the move uses $s^{-1}$, so for graph distance we regard $g\to gs$ and $gs\to g$ as the two orientations of one undirected edge. An involution therefore gives one edge, not two parallel edges. Each edge has length one, regardless of its length in a drawing.^[ref]

A sequence $s_1,\dots ,s_n$ labels the walkIt ends at $h$ exactly when $g^{-1}h=s_1\cdots s_n$. Thus the Cayley graph is a connected graph precisely because $S$ generates $G$, and its graph distance on vertices is the word metric. Left multiplication by $a$ sends the edge from $g$ to $gs$ to the edge from $ag$ to $(ag)s$, preserving its generator label: this is a group action by graph automorphisms and isometries. Viewing groups as symmetries of geometric objects is the motivation of Clay and Margalit's introduction.^[ref]

The useful geometric object includes its allowed moves, not just its drawing. Cayley graphs of the integers show how changing the moves adds shortcuts; word metric balls in the integer lattice show how the same group can acquire visibly different closed balls. Meier's book develops this passage from group operations to graphs as its central motivation.^[ref]

The center of a graph is the set of vertices with minimum graph eccentricity. For a finite connected graph $G$, putHere $d_G$ is shortest-path distance and $\mathrm{rad}G={\min }_{v}e(v)$ is the graph radius. The center solves a facility-placement problem in which the cost is the distance to the farthest vertex.

In a five-vertex path graph, eccentricities are $4,3,2,3,4$, so only the middle vertex is central. In a four-vertex path, the two middle vertices both have eccentricity two. The center of a tree is always a single vertex or a pair of adjacent vertices: repeatedly deleting all leaves exposes its middle. For a general graph the center need not have this form; every vertex of a cycle graph is central.

Minimizing the sum of distances defines a different objective. For example, let $a,b,c,d$ form a path, and attach three additional leaves to $a$. The center is $\{b\}$, but the sum of distances is smallest at $a$. The farthest-distance criterion and the average-distance criterion can therefore select different locations.

A chordal graph is a finite simple graph with no induced cycle of length at least four. Equivalently, every cycle in a graph with at least four vertices has a chord: an edge joining two nonconsecutive vertices of that graph cycle. The word "induced" matters: the induced subgraph on the chosen vertices must contain exactly the cycle's edges, with no chords. Triangles are allowed. This definition concerns adjacency, independently of how the graph is drawn.

For the square cycle graph $C_4$, take vertex set $\{a,b,c,d\}$ and set of edges $\{ab,bc,cd,da\}$, where $ab$ abbreviates $\{a,b\}$. Its graph neighbourhoods are $N(a)=N(c)=\{b,d\}$ and $N(b)=N(d)=\{a,c\}$. The square is an induced cycle, so $C_4$ is not a chordal graph. Add the single edge $ac$: now $N(a)=\{b,c,d\}$, $N(b)=\{a,c\}$, $N(c)=\{a,b,d\}$, and $N(d)=\{a,c\}$. The only possible four-vertex graph cycle has chord $ac$, so this new graph is a chordal graph. Its maximal cliques are exactly $\{a,b,c\}$ and $\{a,c,d\}$.

Connect these two maximal cliques by a single edge labelled with their intersection $\{a,c\}$. The result is a junction tree, also called a clique tree: its nodes are maximal cliques, and the nodes containing any fixed original vertex induce a connected graph inside the tree. Here $a,c$ occur in both nodes, while $b,d$ each occur in just one. This is the running intersection property. More generally, a nonempty connected graph is a chordal graph exactly when its maximal cliques admit such a tree; a disconnected chordal graph can be represented by one such tree per connected component of a graph.^[ref]

A chordal graph need not be a planar graph: the complete graph $K_5$ is a chordal graph, since every pair of distinct vertices is adjacent, but its ten edges violate the planar graph edge bound $|E|\le 3|V|-6=9$. Conversely, even a plane drawing whose bounded faces are triangles can hide an induced cycle: a wheel graph with a four-vertex rim and a central vertex joined to the whole rim still has the rim as an induced cycle. Adding chords to obtain a chordal graph is called triangulation of an undirected graph; this is an adjacency condition, rather than a requirement about planar faces.

Why is this structure useful? The maximal cliques collect local interactions, and the running intersection property makes their overlaps sufficient for consistent communication. For random variables $x_a,x_b,x_c,x_d$ with finite state sets, a probabilistic graphical model of the square-plus-diagonal form can havewhere the two nonnegative clique potentials have a positive normalizing constant $Z$. Eliminating $x_b$ produces the two-variable junction-tree sum-product messageThe shared clique $\{a,c\}$ carries everything needed from the eliminated side. For $q$ possible states per random variable, this example uses two sums over $q^3$ configurations after reusing the message, rather than directly summing over $q^4$ configurations. The example's message is an exact finite sum, with no simulated data. Large maximal cliques still make exact inference expensive; chordal graphs provide a decomposition, not a universal guarantee of cheap inference.^[ref]

For a locally finite connected graph with positive electrical conductances, the number of trees in its wired uniform spanning forest is almost surely a deterministic constant, possibly infinite. A precise characterization uses independent random walks with the conductance transition probabilities. Let $\alpha (w_1,\dots ,w_k)$ be the probability that their full vertex ranges are pairwise disjoint. ThenThe right side depends only on the electrical network. The proof amplifies positive disjointness probability to probabilities arbitrarily close to one, then uses Wilson's algorithm rooted at infinity to obtain at least $k$ components. Conversely, $k$ distinct components force starting vertices admitting positive disjointness probability. This establishes the constant without treating component count as automatically invariant under finite edge changes.^[ref]

An undirected graph is connected when every pair $u,v$ of vertices is joined by a finite graph path. The length-zero graph path consisting of $u$ alone is allowed, so a single-vertex graph is connected. We use the term here for nonempty graphs; conventions for the graph with no vertices vary. Adjacent vertices are connected, but connected vertices need not be adjacent: in our six-vertex example, $A$ reaches $F$ along $A,C,D,F$, although $AF$ is absent.

Allowing walks instead of graph paths gives the same test for a connected graph. Whenever a finite walk revisits a vertex, delete the intervening closed portion; repeating this leaves a graph path with the same endpoints. This simple shortening argument explains why revisiting locations cannot help decide reachability.^[ref]

Connectivity can conceal fragile connections. Removing $CD$ from our graph produces two connected components of a graph, on $\{A,B,C\}$ and $\{D,E,F\}$: $CD$ is a bridge in a graph. Removing $AB$ instead preserves connectivity, because $A,C,B$ supplies an alternative graph path. If edges represent communication links, this calculation says precisely which single link failure separates the two groups. It is a conclusion about this model, independent of the lengths in its drawing.

A tree is a nonempty connected graph without a graph cycle; equivalently, each pair of vertices has exactly one graph path between them. Delete $BC$ and $EF$ from our example. The remaining edgesform a spanning tree: all six vertices remain, with $6-1=5$ edges. The graph path $A,C,D,F$ is now unique, and every remaining edge is a bridge in a graph. Adding $BC$ back creates exactly one graph cycle, $B,A,C,B$. Thus graph cycles express alternative routes, while a tree retains just enough connections for reachability. More generally, a finite connected graph on $n\ge 1$ vertices has at least $n-1$ edges, with equality exactly for a tree. One proof repeatedly removes an edge on a graph cycle, preserving connectivity, until a spanning tree remains.^[ref]

For a finite graph, breadth-first search tests connectivity by starting at one vertex and discovering neighbours in successive layers. Starting at $A$, the layers in our example are $\{A\}$, $\{B,C\}$, $\{D\}$, $\{E,F\}$. Every vertex is discovered, and the layer numbers give graph distances from $A$. Recording one discovery edge for each new vertex constructs a spanning tree. This turns an existence question into both a concrete route and a certificate of connectivity.

Connectivity studies the existence of paths between vertices of a graph and the persistence of paths after vertex or edge deletions. A connected graph is the basic case in which every pair of vertices is linked by a path.

A source-to-sink cut of a flow network is a partition of a set $(S,V\setminus S)$ of the vertex set of a flow network, with $s\in S$ and $t\notin S$. Its cut capacity sums the flow network edge capacities of directed edges leaving $S$.^[ref]

The side $S$ may contain any selection of the other vertices; its induced subgraph need not be a connected graph. For the five-edge flow network with flow network edge capacities $(c_{sa},c_{sb},c_{ab},c_{at},c_{bt})=(3,2,1,2,3)$, $S=\{s\}$ gives the partition of a set $(\{s\},\{a,b,t\})$. Only $s\to a$ and $s\to b$ leave this side, so its cut capacity is $3+2=5$. Choosing $S=\{s,b\}$ instead puts $a\to b$ across the boundary in the incoming direction: its flow network edge capacity is excluded.

The left panel pairs a feasible flow of flow value $5$ with a cut of a flow network of cut capacity $5$, proving that each is optimal. The right panel uses the same feasible flow but a different cut of a flow network: the outward flow is $3+3=6$ and the backward flow is $1$, giving net flow $6-1=5$. That cut capacity is $6$, so it supplies a valid but less sharp bound. The dashed boundary marks a partition of a set, rather than an additional edge or a proposed route.

A cut vertex, also called an articulation point, is a vertex $v$ of an undirected graph whose deletion increases the number of connected components of a graph. Write $k(G)$ for that count, including isolated vertices, and $G-v$ for the graph obtained by removing $v$ and every incident edge. Then the test is $k(G-v)>k(G)$, even when the original graph is disconnected. By contrast, deleting a bridge in a graph removes just one edge and keeps both endpoint vertices. These are different failure models. ^[ref]

First take two triangles in a graph, $ABC$ and $ADE$, sharing only the vertex $A$. This simple graph hasHere $XY$ abbreviates the unordered pair $\{X,Y\}$. Removing $A$ leaves the two connected components of a graph $\{B,C\}$ and $\{D,E\}$, so $k(G-A)=2>1=k(G)$. Removing any other vertex leaves a connected graph. Thus $A$ is the only cut vertex. Every edge lies on a triangle in a graph, so deleting any single edge still leaves a graph path between its endpoints along the other two edges. There are no bridges in a graph.

The generator computes the adjacency lists and checks every vertex deletion, giving the following complete table:

Now take two disjoint triangles in a graph, $ABC$ and $DEF$, joined by the edge $CD$:Deleting only $CD$ preserves all six vertices and leaves connected components of a graph $\{A,B,C\}$ and $\{D,E,F\}$: $CD$ is the only bridge in a graph. Deleting $C$ also removes $AC$ and $BC$, leaving connected components of a graph $\{A,B\}$ and $\{D,E,F\}$. Deleting $D$ instead leaves $\{A,B,C\}$ and $\{E,F\}$. Thus $C,D$ are exactly the cut vertices.

Again the adjacency lists and all vertex deletion counts are computed directly:

For a disconnected example, adjoin an isolated vertex $Z$ to the first graph. Now $k(G)=2$, deleting $A$ gives $k(G-A)=3$, and deleting $Z$ gives $k(G-Z)=1$. The original graph was already disconnected, but $A$ is still a cut vertex; $Z$ is not. The empty graph has count zero, so deleting the only vertex of a one-vertex graph does not create a cut vertex. Nor is either endpoint of a single isolated edge a cut vertex, although that edge is a bridge in a graph.

In a communications network, a cut vertex models a router whose failure separates surviving routers. The shared-$A$ example shows why having an alternative graph path around each individual edge does not protect against failure of the common router. An additional edge from $B$ to $D$ removes this particular vulnerability: after deleting $A$, $B,D$ still join the two sides. A bridge in a graph instead models a vulnerable link. This distinction tells us whether to add another cable between existing devices or a route avoiding an indispensable device; it describes connectivity, rather than traffic capacity or failure probabilities.

The beautiful structural consequence is that a graph splits into biconnected components joined at cut vertices. Within each nontrivial robust piece, a single vertex failure cannot separate the other vertices; between pieces, the shared cut vertices reveal the bottlenecks. Also, a Hamiltonian graph has no cut vertex: deleting any vertex from its Hamilton cycle leaves a Hamiltonian path through all remaining vertices.

Historically, Tarjan's June 1972 paper developed depth-first search as a tool for finding biconnected components in time linear in the numbers of vertices and edges. ^[ref] In his 1983 recollection, Tarjan credited Hopcroft with the original linear-time biconnected component algorithm during their collaboration beginning in summer 1970, and described the published version as a refinement. The historical lesson is how a carefully organized traversal turns many separate deletion experiments into one computation. ^[ref]

For a nontrivial connected graph, return the internal vertices of a rooted tree in a depth-first search tree. Undirected depth-first search has only ancestor-descendant edges, so every edge meets an internal vertex; the internal vertices are connected along parent paths. For a maximum-size proper matching in a rooted tree $M$ of that tree, existence of which is assumed, $|I|=|V|-|L|\le |V|-|U(M)|=2|M|$. Every vertex cover has size at least $|M|$, proving the factor two. The implementation needs only depth-first search, hence linear time complexity; it does not compute the analysis matching in a graph.

Depth-first search is a graph search that follows a branch as far as it can before returning to try the next unexplored branch. Starting at a vertex $s$ of an undirected graph on a finite set of vertices, mark $s$ as discovered. Scan its graph neighbours in a fixed order: for each undiscovered vertex, record the current vertex as its parent in a rooted tree in the discovery tree, mark the new vertex, and explore it immediately using recursion. Return when its adjacency list is exhausted. Discovery marks persist after return, preventing a graph cycle from causing repeated exploration. This branch-and-return rule is the maze exploration described in MIT's notes. ^[ref]

Use the same unweighted undirected graph as breadth-first search, with vertex set $V=\{A,B,C,D,E,F,G\}$ and edgeswhere $XY$ denotes the unordered pair $\{X,Y\}$. Visit graph neighbours in lexicographic order of their one-letter labels, starting at $A$. The discovery order, computed by the diagram's generator, isThe discovery edges in order are $AB,BD,BE,EC,CF,FG$. They form a rooted tree spanning the reached vertices, with root of a tree $A$. The other edges, $AC$ and $EG$, remain in the original graph but never discover a new vertex. In particular, the graph cycle $A,B,E,C,A$ causes no repeated discovery.

At $D$ there is no undiscovered graph neighbour, so its call returns to $B$. The next call enters $E$, then $C$, then $F$, then $G$. Immediately after discovering $G$, the call stack is $A,B,E,C,F,G$, with $G$ on top; $D$ is still discovered but its call has finished. The maximum call stack depth is six, although seven vertices are discovered. Following the recorded parent in a rooted tree backwards at each step reconstructs a graph path; here the route from $A$ to $G$ is $A,B,E,C,F,G$, using five edges. Breadth-first search instead finds $A,B,E,G$, using three edges. Depth-first search establishes reachability but does not guarantee the fewest edges.

An explicit LIFO stack can replace recursion, but its entries and discovery rule must be specified. To reproduce this exact traversal, store one frame per active vertex, retaining the next position in its adjacency list. Advance the top frame by one graph neighbour; if that vertex is undiscovered, mark it and push its frame immediately; pop an exhausted frame. Pushing all undiscovered graph neighbours at once and marking them on push is a different iterative variant. Even pushing them in reverse lexicographic order does not generally reproduce the parent in a rooted tree relationships above: $C$ would already be marked while processing $A$, so $E$ could not discover $C$.

With an adjacency list representation, constant-cost discovery checks and frame operations, and the graph neighbours already ordered, depth-first search has time complexity $O(|V|+|E|)$ in Big O notation. Each reached vertex is entered once and each incident edge is examined at most twice in an undirected graph; initializing all discovery marks costs $O(|V|)$. The call stack has depth at most $|V|$, since no active vertex repeats. Discovery marks and the recorded parent in a rooted tree for each nonroot vertex also need $O(|V|)$ auxiliary space complexity, excluding storage for the graph. Sorting adjacency lists to obtain a chosen order is additional work. ^[ref]

One start explores only its connected component of a graph. To cover a disconnected graph, start again at every still-undiscovered vertex while keeping all discovery marks; the discovery trees form a forest. Depth-first search is useful for reachability, constructing a spanning tree of a connected graph, and organizing the return steps needed by a depth-first traversal of a tree. For detecting a graph cycle in a simple graph that is an undirected graph, an edge to an already discovered vertex other than the current parent in a rooted tree closes a graph cycle; the immediate return edge to that parent in a rooted tree is not itself a graph cycle.

// Target: computer-science.bigb

An ideal differential measurement constrains two sky temperatures through $d_{ij}=T_i-T_j$. Represent every observed pair by a row of a matrix $A$, with $+1$ in column $i$, $-1$ in column $j$ and zeros elsewhere. Then $d=AT$ and $A1=0$: adding any constant to the entire map changes no measurement. This null space is the angular angular monopole ambiguity, explaining why a directional comparison alone cannot measure the absolute blackbody temperature established by FIRAS.

For an original three-pixel example choose $T=(30,-10,-20)\mu \mathrm{K}$ relative to an unspecified mean. Measurements $d_{12}=40\mu \mathrm{K}$ and $d_{23}=10\mu \mathrm{K}$ determine the third difference $d_{31}=-50\mu \mathrm{K}$. The zero-mean condition recovers the chosen map, but adding $100\mu \mathrm{K}$ to every pixel gives identical differences. The closure sum $d_{12}+d_{23}+d_{31}=0$ checks consistency without establishing the mean.

More generally, the graph joining measured pixel pairs must be a connected graph. If $AT=0$, each edge forces equal endpoint temperatures; a path propagates that equality throughout each connected component. Thus a connected observation graph leaves exactly one undetermined constant, whereas disconnected coverage leaves one per component. Fixing the mean makes an ideal connected reconstruction unique; real noise calls for least-squares solution fitting and uncertainty analysis.

The figure chooses a cosmic microwave background quadrupole sky $T(\theta )=Q{P}_2(\cos \theta )$, with $Q=30\mu \mathrm{K}$ and Legendre polynomial $P_2(u)=(3u^2-1)/2$. On a great-circle scan, with $\varphi$ measured from the chosen symmetry axis and ideal beams separated by $\alpha =60^{\circ }$,Expanding ${\cos }^2\varphi =(1+\cos 2\varphi )/2$ proves the formula; its peak magnitude is $45\sqrt{3}/2\simeq 38.9711\mu \mathrm{K}$. A constant sky gives zero. This deliberately restricted scan also misses some angular patterns, so rejecting the angular monopole is not a proof of complete sky reconstruction. Broad, overlapping coverage supplies the missing constraints; finite antenna beam patterns additionally smooth small angular features.

The continuous-time diffusion model on a finite undirected graph uses the Graph Laplacian in the ordinary differential equationAcross each edge, the higher value sends an amount proportional to the difference to the lower value. The two contributions cancel in the sum over vertices, giving conservation of mass:If the entries represent masses, start with nonnegative entries; signed entries can instead describe deviations from a common level. For the three-vertex path graph, the initial vector $x(0)=(1,0,-1{)}^T$ is an eigenvector with eigenvalue one, so the exact solution isMore generally, an eigenvector with eigenvalue $\lambda$ decays by $e^{-\lambda t}$. The spectral theorem for real symmetric matrices supplies an orthonormal basis of these eigenvectors, and the kernel of a graph Laplacian identifies the surviving terms. On a nonempty connected graph,The limiting value is the original arithmetic mean. On a disconnected graph, each connected component of a graph converges to its own original arithmetic mean. The same matrix therefore measures neighbour disagreement, detects disconnected regions and governs relaxation toward agreement.