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Every doubly stochastic matrix is a convex combination of permutation matrices. Equivalently, the extreme points of the doubly stochastic convex polytope are exactly the permutation matrices. The extreme-point assertion follows from an alternating-cycle perturbation of a doubly stochastic matrix; the convex-hull assertion follows because a bounded finite-dimensional polytope is the convex hull of its vertices. Permutation matrices are extreme since their zero entries force the same zeros in every convex decomposition.

For an integer $0\le k\le n$, the capped simplexis the convex hull of the zero-one vectors having exactly $k$ entries equal to one. It is a convex polytope whose extreme points are precisely these vectors.

If a feasible vector has a fractional coordinate, the integer coordinate sum forces at least two fractional coordinates. Perturbing one upward and the other downward by a sufficiently small amount gives two distinct feasible vectors whose midpoint is the original vector. It is therefore not an extreme point. Conversely, a zero-one vector cannot be a nontrivial convex combination of points of $[0,1{]}^n$, since each coordinate is already at a bound.

A convex polygon is a two-dimensional convex polytope, equivalently the convex hull of a finite planar set with nonempty interior. Its vertices are the extreme points of that hull. Saying that given points form a convex polygon means every specified point is a vertex; points inside the hull or in the interior of a straight boundary segment do not count as vertices.

A facet of a full-dimensional convex polytope in ${\mathbb{R}}^n$ is a face of dimension $n-1$. Each irredundant supporting closed half-space defines one facet.

If $r\ge \sqrt{2}$ and a convex polytope satisfies $r^{-1}{B}^n\subseteq P\subseteq B^n$, then it has at least $(1-r^{-2}{)}^{-n/2}\ge e^{n/(2r^2)}$ facets. Its outward facet normals define spherical caps covering the unit sphere; the spherical cap area upper bound supplies the estimate.

For a finite matroid $M=(E,\mathcal{I})$, the convex hull of the indicator vectors of its matroid independent sets is exactly the displayed system, where $x(A)={\sum }_{e\in A}{x}_e$. It is an integral polytope. The equality $x(E)=r_M(E)$ instead gives the face corresponding to matroid bases; independence does not require that equality.

Proof. If $I$ is independent, $|I\cap A|\le r_M(A)$, so every independent indicator and every convex combination of them satisfies the inequalities. For the reverse inclusion, take an arbitrary real objective $c$. List its positive coordinates in nonincreasing order as $c_1\ge \cdots \ge c_k>0$, put $c_{k+1}=0$ and let $S_i$ be the first $i$ elements in this ordering. Run the matroid greedy algorithm on these elements, obtaining an independent set $I$. At every prefix the selected elements are maximal in $S_i$: an earlier rejected element cannot become admissible after further additions, by heredity. The augmentation axiom implies that every maximal independent set in $S_i$ has size $r_M(S_i)$. Hence $|I\cap S_i|=r_M(S_i)$.

For every point satisfying the displayed system, nonnegativity and telescoping giveThus every linear objective is maximized by an independent indicator, with the empty set used if all coefficients are nonpositive. Singleton rank inequalities bound every coordinate between zero and one, making the system compact. If it contained a point outside the compact convex hull of independent indicators, the finite-dimensional separating hyperplane theorem would supply an objective strictly larger at that point than at every independent indicator, a contradiction. This proves equality and integrality without assuming the desired convex-hull description in advance. Negative objective coefficients are discarded explicitly; maximizing over independent sets differs from maximizing over bases when such coefficients occur.

Assume $n\ge 1$ and the integer coefficient bound $U\ge 1$. There are $m(n+1)$ integers to record. Each signed integer in $[-U,U]$ can be encoded in $\lceil {\log }_2(2U+1)\rceil$ bits, so a fixed-width dense encoding needsbits, including the dimensions and bound. Its size is $\boxed{O(mn\log (U+1))}$. This is the bit-input model used for polynomial time; an arithmetic-operation count alone would not suffice.

We first construct a rational feasibility certificate for integer inequalities. Choose a feasible point and an orthant containing it, adding coordinate inequalities of the form $x_i\ge 0$ or $-x_i\ge 0$. The intersection $Q$ is nonempty. Starting at a point of $Q$, if the active constraint normals have rank less than $n$, choose a nonzero direction $d$ annihilating them. A sufficiently short move in either direction remains feasible, since every inactive constraint has positive slack. At least one direction has a finite endpoint: a nonzero coordinate of $d$ can be oriented toward the orthant boundary. Move to the first constraint met along such a direction. Its normal does not annihilate $d$, so it is independent of the previously active normals. Their rank strictly increases. In at most $n$ moves we obtain a point with $n$ independent active constraints.

At that point $Bx=d_0$ for an invertible integer $n\times n$ matrix $B$. Its rows are original constraint rows or signed coordinate rows, and $d_0$ has entries from $b$ or zero. All entries have absolute value at most $U$. By Cramer's rule,where $B_i$ replaces column $i$ by $d_0$. The Hadamard determinant inequality bounds both determinants by $(\sqrt{n}U{)}^n\le (nU{)}^n$. The denominator is a nonzero integer. Hence each coordinate has the claimed representation with numerator and denominator bounded in absolute value by $R=(nU{)}^n$. The orthant argument is needed because $P$ itself may contain lines and have no vertex. If $U=0$ is allowed literally, all coefficients and right sides vanish, and $x=0$ is feasible; the printed fraction bound then needs $U$ replaced by $\max (1,U)$, since a nonzero denominator cannot have absolute value at most zero. The bound just proved uses the customary positive coefficient bound.

Recording $n$ such fractions takes $O(n^2\log (nU))$ bits. A verifier checks that denominators are nonzero, changes their signs if needed, and verifies all inequalities by exact integer arithmetic after multiplying by positive denominators. The intermediate integers still have polynomial bit length. Thus a yes-instance has a polynomial-size, polynomial-time verifiable witness, and $\boxed{\mathrm{FP}\in \mathbf{NP}}$.

For optimization, let $c$ be rational with a finite binary encoding; clear its denominators to make it integer, which multiplies the objective by a positive constant and preserves its minimizers. All bit bounds below include the encoding length of $c$. Let $V=\max (1,{\max }_i|c_i|)$ and $B_0=nVR$.

A useful refinement of the preceding vertex argument is that if an objective is bounded below, a feasible point can be moved to an orthant vertex without increasing its objective. Choose the annihilating direction to have $c^{T}d\le 0$. If the inequality is strict, the forward feasible interval must have a finite endpoint, since otherwise the objective would be unbounded below. If $c^{T}d=0$, choose either finite endpoint supplied by the orthant. Rank increases as before. There are only finitely many such vertices across the finitely many orthants. Every feasible point therefore has one of these vertices with no larger objective; taking the least vertex value proves that a finite optimum is attained at one of them.

The Cramer representations at this optimum have one common denominator $\det B$, of magnitude at most $R$. Therefore its objective value $z^{\ast }$ has a representation with denominator at most $R$ and numerator at most $B_0$. In particular $|z^{\ast }|\le B_0$. No enumeration of orthants or vertices is part of the algorithm; they provide only polynomial-bit bounds.

Here is an exact linear optimization from a feasibility oracle. First call the assumed ellipsoid method feasibility routine on $P$; if it is empty, report infeasibility. Otherwise query feasibility with the extra inequality $c^{T}x\le -B_0-1$. A yes answer proves unboundedness below, because any finite optimum would be at least $-B_0$; an unbounded objective necessarily gives a yes answer. Thus the infeasible, unbounded and finite-optimum cases are distinguished using two feasibility calls.

In the finite case, maintain an infeasible lower threshold $-B_0-1$ and a feasible upper threshold $B_0$. Bisect the interval, querying $c^{T}x\le t$ each time. After polynomially many calls its length is less than $1/(2R^2)$, and $z^{\ast }$ lies in the resulting interval. Distinct reduced fractions with positive denominators at most $R$ differ by at least $1/R^2$, so there is only one possible value. It can be recovered in polynomial time by continued fractions: for the midpoint $t$, $|z^{\ast }-t|<1/(4R^2)\le 1/(2q^2)$, where $q\le R$ is the denominator of $z^{\ast }$. Thus $z^{\ast }$ occurs among the continued-fraction convergents of $t$. Compute those by the Euclidean algorithm and choose the unique fraction with denominator at most $R$ in the interval. If $t=z^{\ast }$, use the final exact convergent. The Euclidean algorithm and all thresholds have polynomial bit length.

Even if the feasibility routine returns only a decision, an optimal point can be recovered without assuming a witness-producing oracle. Add $c^{T}x=z^{\ast }$ and $-R\le x_i\le R$ for all $i$. This bounded optimal convex polytope is nonempty, since the optimal vertex above lies in it. Clear the denominator of $z^{\ast }$ and setAll vertex coordinates of this fixed convex polytope have denominators at most $D_0$ by the same determinant argument. Minimize $x_1$, recover its exact minimum by feasibility bisection and rational reconstruction, then fix it; do the same for $x_2$, and so on. Each restriction is an optimal face of the same original bounded convex polytope, so its vertices are original vertices. The single denominator bound $D_0$ therefore works at every step; repeated new precision bounds do not grow exponentially. After $n$ steps all coordinates are fixed to a feasible optimizer. There are polynomially many calls on polynomial-length rational instances, converted to integer inequalities by clearing denominators. Consequently linear optimization has a polynomial-time algorithm whenever linear feasibility does, with exact optimal value and optimizer, and correct reports of infeasibility or unboundedness.

The equilibrium oddness theorem gives a finite odd number of Nash equilibria in a finite nondegenerate bimatrix game. One way to see the parity is through the endpoint argument underlying the Lemke-Howson algorithm. Add sufficiently large constants to each payoff matrix so all entries are positive; this does not change best responses. Form the bounded convex polytopesLabel an $s_i=0$ facet by row label $i$ and a tight $(Q^{T}s{)}_j=1$ facet by column label $n+j$; in $Y$, label a tight $(Pt{)}_i=1$ facet by $i$ and $t_j=0$ by $n+j$. A completely labelled vertex of a polytope pair has either both vectors zero, or both nonzero and normalizes to a Nash equilibrium. In the latter case the labels express exactly the equilibrium best-response and zero-probability conditions. Nondegeneracy makes each vertex of a polytope have exactly $n$ distinct incident labels and each equilibrium correspond to one such pair.

Fix a label to drop and retain all pairs carrying every other label. Include the product-convex polytope edges that retain those labels. At a completely labelled pair there is one possible outgoing edge, obtained by dropping the fixed label. At any other retained vertex of a polytope, that label is missing and one other label is duplicated; dropping either copy gives the two incident edges. Thus this finite graph consists of paths and cycles, with its endpoints precisely the completely labelled pairs. A finite graph has an even number of degree-one vertices. Since one endpoint is the artificial zero pair, the number of genuine equilibrium endpoints is odd. This proves the required equilibrium oddness theorem and finiteness, without claiming that every equilibrium is reached from the zero pair on the same path.

For a symmetric bimatrix game, swapping players maps $(s,t)$ to $(t,s)$. Every nonsymmetric equilibrium lies in a distinct two-element pair; the fixed points of this involution are exactly the symmetric equilibria. Removing even-sized pairs from an odd total leaves an odd number of symmetric equilibria. In the present game, two equilibria form the swapped pair and the remaining one is $(x,x)$ with $x=(1/2,0,1/2)$. There is exactly one symmetric equilibrium, as required.

Let $m$ be the number of facets of the convex polytope $P$. Its facet description can be writtenBecause $r^{-1}{B}^n\subset P$, each supporting closed half-space has $h_j\ge 1/r$. For any $\theta \in S^{n-1}$, follow the ray $t\theta$ to its last point in $P$. Its radius $t_{\theta }$ is at most one because $P\subset B^n$. At least one facet is active there, givingThus the unit sphere is covered by the $m$ spherical caps centred at $u_j$ with angular radius $\rho =\arccos (1/r)$. Since $r\ge \sqrt{2}$, this radius lies in $[\pi /4,\pi /2)$ and the spherical cap area upper bound applies. Subadditivity of surface area givesUsing $\log (1-x)\le -x$ for $0<x<1$,The facet lower bound for ball approximations is exponential in the dimension.

Choose a maximal subset $U\subset S^{n-1}$ whose distinct points have Euclidean distance greater than one. The volumetric bound for Euclidean metric nets makes the construction finite: the open Euclidean balls of radius $1/2$ about the points of $U$ are disjoint and all lie in $(3/2)B^n$, soMaximality means that $U$ is a metric net of radius one on the unit sphere. Thus every $\theta \in S^{n-1}$ has some $u\in U$ with $\Vert u-\theta \Vert \le 1$. Since both are unit vectors,Now intersect the corresponding closed half-spaces:The Cauchy-Schwarz inequality shows that $(1/2)B^n\subset P$. Conversely, write any nonzero $x\in P$ as $x=t\theta$ and choose $u$ as above. Then $t/2\le t(u\cdot \theta )=u\cdot x\le 1/2$, so $t\le 1$. Hence $P\subset B^n$, which also proves boundedness. This convex polytope has at most $|U|\le 3^n$ facets, since redundant inequalities can only reduce their number. It meets the required inclusions with $r=2$.

Suppose $y$ in the capped simplex has a coordinate strictly between zero and one. Since ${\sum }_i{y}_i=k$ is an integer, it cannot have exactly one such coordinate: the other coordinates would contribute an integer sum. Hence there are distinct $i,j$ with $0<y_i,y_j<1$.

Choose $0<\epsilon <\min \{y_i,1-y_i,y_j,1-y_j\}$ and set $y^{\pm }=y\pm \epsilon (e_i-e_j)$, where $e_i,e_j$ are standard basis vectors. Both perturbed vectors satisfy the coordinate bounds and have the same coordinate sum. They are distinct and $y=(y^{+}+y^{-})/2$, so $y$ is not an extreme point.

Conversely, every zero-one vector in this convex polytope is an extreme point. If it were a nontrivial convex combination of two feasible vectors, each coordinate equal to zero would force both corresponding coordinates to be zero, and each coordinate equal to one would force both to be one. Thus the two vectors would equal the original one. The extreme points are therefore exactly the indicators of $k$-element subsets.

The origin belongs to $S$, so $S$ is nonempty. Each displayed inequality defines a closed half-space, and every half-space is a convex set. An intersection of convex sets is convex: if $x,y\in S$ and $0\le t\le 1$, each linear inequality remains valid for $(1-t)x+ty$. Thus $S$ is a nonempty convex polytope.

Orient every edge of the feasible convex polytope in the direction of increasing objective. WriteThe possible improving simplex method paths from $O$ to $A$ areThese exhaust the directed edge graph: the ordinary cube edges remain except those incident to the removed vertex $111$; each such edge ends at one of $A,B,C$; and the cutting face contributes the triangle with edges $AB,AC,BC$.

The first two paths use three pivots, while the two paths through both $C$ and $B$ use five. Therefore the smallest and largest possible numbers of simplex steps areand the total number of distinct outcomes isThis is the simplex paths on a cube with one truncated corner calculation.

Take $a=1$ and $b=k$, sois the capped simplex. A linear objective over this convex polytope attains its maximum at a zero-one extreme point. Choosing the $k$ coordinates at which $x_i$ is largest givesEquivalently, an exchange of weight from a smaller component to a larger one never decreases the objective. Thus the sum of the largest components is the support function ${\sigma }_C$.

Part c now givesBy the projection onto a box-constrained hyperplane, $v=P_C(y/t)$ hasConsequently the proximal operator is evaluated by solving this one-dimensional equation for $\nu$, then substituting the resulting projection.

For a linear polyhedron $P=\{x\in {\mathbb{R}}^n:Ax\le b\}$, a vertex of a linear polyhedron is a feasible point $p$ at which the normals of the active constraints span ${\mathbb{R}}^n$. Equivalently, $n$ linearly independent active constraints uniquely determine $p$. A linear polyhedron may be unbounded or lower-dimensional; the definition applies in those cases as well as to a convex polytope.

Such a point is an extreme point. Indeed, if $p=\lambda u+(1-\lambda )v$ with $u,v\in P$ and $0<\lambda <1$, every active constraint at $p$ must also be an equality at $u$ and $v$. Selecting $n$ independent active rows forces $u=v=p$. Conversely, if the active normals do not span ${\mathbb{R}}^n$, take a nonzero vector $h$ annihilated by all active rows. The finitely many inactive inequalities have positive slack, so both $p+\epsilon h$ and $p-\epsilon h$ remain feasible for sufficiently small $\epsilon >0$. Their midpoint is $p$, so $p$ is not an extreme point. Thus these notions coincide for a linear polyhedron.

A vertex of a convex polytope is an extreme point: it cannot be expressed as a nontrivial convex combination of two distinct points in that polytope. In a full-dimensional simple polytope of dimension $n$, a vertex lies on exactly $n$ incident facets.