graph partitioning algorithms and laplacian eigenvalues · graph partitioning algorithms and...

Graph Partitioning Algorithms and LaplacianEigenvalues

Luca TrevisanStanford

Based on work with Tsz Chiu Kwok, Lap Chi Lau, James Lee,Yin Tat Lee, and Shayan Oveis Gharan

spectral graph theory

Use linear algebra to

• understand/characterize graph properties

• design efficient graph algorithms

linear algebra and graphs

Take an undirected graph G = (V ,E ), consider matrix

A[i , j ] := weight of edge (i , j)

eigenvalues −→ combinatorial propertieseigenvenctors −→ algorithms for combinatorial problems

linear algebra and graphs

Take an undirected graph G = (V ,E ), consider matrix

A[i , j ] := weight of edge (i , j)

eigenvalues −→ combinatorial propertieseigenvenctors −→ algorithms for combinatorial problems

linear algebra and graphs

Take an undirected graph G = (V ,E ), consider matrix

A[i , j ] := weight of edge (i , j)

eigenvalues −→ combinatorial propertieseigenvenctors −→ algorithms for combinatorial problems

fundamental thm of spectral graph theory

G = (V ,E ) undirected, A adjacency matrix, dv degree of v , Ddiagonal matrix of degrees

L := I − D−12AD−

12

L is symmetric, all its eigenvalues are real and

• 0 = λ1 ≤ λ2 ≤ · · · ≤ λn ≤ 2

• λk = 0 iff G has ≥ k connected components

• λn = 2 iff G has a bipartite connected component

fundamental thm of spectral graph theory

G = (V ,E ) undirected, A adjacency matrix, dv degree of v , Ddiagonal matrix of degrees

L := I − D−12AD−

12

L is symmetric, all its eigenvalues are real and

• 0 = λ1 ≤ λ2 ≤ · · · ≤ λn ≤ 2

• λk = 0 iff G has ≥ k connected components

• λn = 2 iff G has a bipartite connected component

fundamental thm of spectral graph theory

G = (V ,E ) undirected, A adjacency matrix, dv degree of v , Ddiagonal matrix of degrees

L := I − D−12AD−

12

L is symmetric, all its eigenvalues are real and

• 0 = λ1 ≤ λ2 ≤ · · · ≤ λn ≤ 2

• λk = 0 iff G has ≥ k connected components

• λn = 2 iff G has a bipartite connected component

fundamental thm of spectral graph theory

G = (V ,E ) undirected, A adjacency matrix, dv degree of v , Ddiagonal matrix of degrees

L := I − D−12AD−

12

L is symmetric, all its eigenvalues are real and

• 0 = λ1 ≤ λ2 ≤ · · · ≤ λn ≤ 2

• λk = 0 iff G has ≥ k connected components

• λn = 2 iff G has a bipartite connected component

0,2

3,

2

3,

2

3,

2

3,

2

3,

5

3,

5

3,

5

3,

5

3

0,2

3,

2

3,

2

3,

2

3,

2

3,

5

3,

5

3,

5

3,

5

3

0, 0, .69, .69, 1.5, 1.5, 1.8, 1.8

0, 0, .69, .69, 1.5, 1.5, 1.8, 1.8

0,2

3,

2

3,

2

3,

4

3,

4

3,

4

3, 2

0,2

3,

2

3,

2

3,

4

3,

4

3,

4

3, 2

eigenvalues as solutions to optimizationproblems

If M ∈ Rn×n is symmetric, and

λ1 ≤ λ2 ≤ · · · ≤ λnare eigenvalues counted with multiplicities, then

λk = minA k−dim subspace of RV

maxx∈A

xTMx

xT x

and eigenvectors of λ1, . . . , λk are basis of opt A

why eigenvalues relate to connectedcomponents

If G = (V ,E ) is undirected and L is Laplacian, then

xTLx

xT x=

∑(u,v)∈E |xu − xv |2∑

v dv · x2v

If λ1 ≤ λ2 ≤ · · ·λn are eigenvalues of L with multiplicities,

λk = minA k−dim subspace of RV

maxx∈A

∑(u,v)∈E |xu − xv |2∑

v dv · x2v

Note:∑

(u,v)∈E |xu − xv |2 = 0 iff x constant on connectedcomponents

why eigenvalues relate to connectedcomponents

If G = (V ,E ) is undirected and L is Laplacian, then

xTLx

xT x=

∑(u,v)∈E |xu − xv |2∑

v dv · x2v

If λ1 ≤ λ2 ≤ · · ·λn are eigenvalues of L with multiplicities,

λk = minA k−dim subspace of RV

maxx∈A

∑(u,v)∈E |xu − xv |2∑

v dv · x2v

Note:∑

(u,v)∈E |xu − xv |2 = 0 iff x constant on connectedcomponents

why eigenvalues relate to connectedcomponents

If G = (V ,E ) is undirected and L is Laplacian, then

xTLx

xT x=

∑(u,v)∈E |xu − xv |2∑

v dv · x2v

If λ1 ≤ λ2 ≤ · · ·λn are eigenvalues of L with multiplicities,

λk = minA k−dim subspace of RV

maxx∈A

∑(u,v)∈E |xu − xv |2∑

v dv · x2v

Note:∑

(u,v)∈E |xu − xv |2 = 0 iff x constant on connectedcomponents

if G has ≥ k connected components

λk = minA k−dim subspace of RV

maxx∈A

∑(u,v)∈E |xu − xv |2∑

v dv · x2v

Consider the space of all vectors that are constant in eachconnected component; it has dimension at least k and so itwitnesses λk = 0.

if G has ≥ k connected components

λk = minA k−dim subspace of RV

maxx∈A

∑(u,v)∈E |xu − xv |2∑

v dv · x2v

Consider the space of all vectors that are constant in eachconnected component; it has dimension at least k and so itwitnesses λk = 0.

if λk = 0

λk = minA k−dim subspace of RV

maxx∈A

∑(u,v)∈E |xu − xv |2∑

v dv · x2v

The eigenvectors x(1), . . . , x(k) of λ1, . . . , λk are all constant onconnected components.The mapping

F (v) := (x(1)v , x

(2)v , . . . , x

(k)v )

is constant on connected components

It maps to at least k distinct points

G has at least k connected components

if λk = 0

λk = minA k−dim subspace of RV

maxx∈A

∑(u,v)∈E |xu − xv |2∑

v dv · x2v

The eigenvectors x(1), . . . , x(k) of λ1, . . . , λk are all constant onconnected components.

The mapping

F (v) := (x(1)v , x

(2)v , . . . , x

(k)v )

is constant on connected components

It maps to at least k distinct points

G has at least k connected components

if λk = 0

λk = minA k−dim subspace of RV

maxx∈A

∑(u,v)∈E |xu − xv |2∑

v dv · x2v

The eigenvectors x(1), . . . , x(k) of λ1, . . . , λk are all constant onconnected components.The mapping

F (v) := (x(1)v , x

(2)v , . . . , x

(k)v )

is constant on connected components

It maps to at least k distinct points

G has at least k connected components

if λk = 0

λk = minA k−dim subspace of RV

maxx∈A

∑(u,v)∈E |xu − xv |2∑

v dv · x2v

The eigenvectors x(1), . . . , x(k) of λ1, . . . , λk are all constant onconnected components.The mapping

F (v) := (x(1)v , x

(2)v , . . . , x

(k)v )

is constant on connected components

It maps to at least k distinct points

G has at least k connected components

if λk = 0

λk = minA k−dim subspace of RV

maxx∈A

∑(u,v)∈E |xu − xv |2∑

v dv · x2v

The eigenvectors x(1), . . . , x(k) of λ1, . . . , λk are all constant onconnected components.The mapping

F (v) := (x(1)v , x

(2)v , . . . , x

(k)v )

is constant on connected components

It maps to at least k distinct points

G has at least k connected components

conductance

vol(S) :=∑

v∈S dv

φ(S) :=E (S , S)

vol(S)

φ(G ) := minS⊆V :0<vol(S)≤ 1

2vol(V )

φ(S)

In regular graphs, same as edge expansion and, up to factor of 2non-uniform sparsest cut

conductance

0, 0, .69, .69, 1.5, 1.5, 1.8, 1.8

φ(G ) = φ(0, 1, 2) =0

6= 0

conductance

0, .055, .055, .211, . . .

φ(G ) =2

18

conductance

0,2

3,

2

3,

2

3,

2

3,

2

3,

5

3,

5

3,

5

3,

5

3

φ(G ) =5

15

finding S of small conductance

• G is a social network

• S is a set of users more likely to be friends with each otherthan anyone else

• G is a similarity graph on items of a data sets• S are items more similar to each other than to other items• [Shi, Malik]: image segmentation

• G is an input of a NP-hard optimization problem• Recurse on S , V − S , use dynamic programming to combine in

time 2O(|E(S,S)|)

finding S of small conductance

• G is a social network• S is a set of users more likely to be friends with each other

than anyone else

• G is a similarity graph on items of a data sets• S are items more similar to each other than to other items• [Shi, Malik]: image segmentation

• G is an input of a NP-hard optimization problem• Recurse on S , V − S , use dynamic programming to combine in

time 2O(|E(S,S)|)

finding S of small conductance

• G is a social network• S is a set of users more likely to be friends with each other

than anyone else

• G is a similarity graph on items of a data sets

• S are items more similar to each other than to other items• [Shi, Malik]: image segmentation

• G is an input of a NP-hard optimization problem• Recurse on S , V − S , use dynamic programming to combine in

time 2O(|E(S,S)|)

finding S of small conductance

• G is a social network• S is a set of users more likely to be friends with each other

than anyone else

• G is a similarity graph on items of a data sets• S are items more similar to each other than to other items• [Shi, Malik]: image segmentation

• G is an input of a NP-hard optimization problem• Recurse on S , V − S , use dynamic programming to combine in

time 2O(|E(S,S)|)

finding S of small conductance

• G is a social network• S is a set of users more likely to be friends with each other

than anyone else

• G is a similarity graph on items of a data sets• S are items more similar to each other than to other items• [Shi, Malik]: image segmentation

• G is an input of a NP-hard optimization problem• Recurse on S , V − S , use dynamic programming to combine in

time 2O(|E(S,S)|)

conductance versus λ2

Recall:λ2 = 0⇔ G disconnected

⇔ φ(G ) = 0

Cheeger inequality (Alon, Milman, Dodzuik, mid-1980s):

λ2

2≤ φ(G ) ≤

√2λ2

conductance versus λ2

Recall:λ2 = 0⇔ G disconnected⇔ φ(G ) = 0

Cheeger inequality (Alon, Milman, Dodzuik, mid-1980s):

λ2

2≤ φ(G ) ≤

√2λ2

conductance versus λ2

Recall:λ2 = 0⇔ G disconnected⇔ φ(G ) = 0

Cheeger inequality (Alon, Milman, Dodzuik, mid-1980s):

λ2

2≤ φ(G ) ≤

√2λ2

Cheeger inequality

λ2

2≤ φ(G ) ≤

√2λ2

Constructive proof of φ(G ) ≤√

2λ2:

• given eigenvector x of λ2.

• sort vertices v according to xv

• a set of vertices S occurring as initial segment or finalsegment of sorted order has

φ(S) ≤√

2λ2 ≤ 2√φ(G )

sweep algorithm can be implemented in O(|V |+ |E |) time

Cheeger inequality

λ2

2≤ φ(G ) ≤

√2λ2

Constructive proof of φ(G ) ≤√

2λ2:

• given eigenvector x of λ2.

• sort vertices v according to xv

• a set of vertices S occurring as initial segment or finalsegment of sorted order has

φ(S) ≤√

2λ2

≤ 2√φ(G )

sweep algorithm can be implemented in O(|V |+ |E |) time

Cheeger inequality

λ2

2≤ φ(G ) ≤

√2λ2

Constructive proof of φ(G ) ≤√

2λ2:

• given eigenvector x of λ2.

• sort vertices v according to xv

• a set of vertices S occurring as initial segment or finalsegment of sorted order has

φ(S) ≤√

2λ2 ≤ 2√φ(G )

sweep algorithm can be implemented in O(|V |+ |E |) time

applications

λ2

2≤ φ(G ) ≤

√2λ2

• rigorous analysis of sweep algorithm(practical performance usually better than worst-case analysis)

• to certify φ ≥ Ω(1), enough to certify λ2 ≥ Ω(1)

• mixing time of random walk is O(

1λ2

log |V |)

and so also

O

(1

φ2(G )log |V |

)

applications

λ2

2≤ φ(G ) ≤

√2λ2

• rigorous analysis of sweep algorithm(practical performance usually better than worst-case analysis)

• to certify φ ≥ Ω(1), enough to certify λ2 ≥ Ω(1)

• mixing time of random walk is O(

1λ2

log |V |)

and so also

O

(1

φ2(G )log |V |

)

applications

λ2

2≤ φ(G ) ≤

√2λ2

• rigorous analysis of sweep algorithm(practical performance usually better than worst-case analysis)

• to certify φ ≥ Ω(1), enough to certify λ2 ≥ Ω(1)

• mixing time of random walk is O(

1λ2

log |V |)

and so also

O

(1

φ2(G )log |V |

)

applications

λ2

2≤ φ(G ) ≤

√2λ2

• rigorous analysis of sweep algorithm(practical performance usually better than worst-case analysis)

• to certify φ ≥ Ω(1), enough to certify λ2 ≥ Ω(1)

• mixing time of random walk is O(

1λ2

log |V |)

and so also

O

(1

φ2(G )log |V |

)

Lee, Oveis-Gharan, T, 2012

From fundamental theorem of spectral graph theory:

• λk = 0⇔ G has ≥ k connected components

We prove:

• λk small ⇔ G has ≥ k disjoint sets of small conductance

Lee, Oveis-Gharan, T, 2012

From fundamental theorem of spectral graph theory:

• λk = 0⇔ G has ≥ k connected components

We prove:

• λk small ⇔ G has ≥ k disjoint sets of small conductance

order-k conductance

Define order-k conductance

φk(G ) := minS1,...,Sk disjoint

maxi=1,...,k

φ(Si )

Note:

• φ(G ) = φ2(G )

• φk(G ) = 0⇔ G has ≥ k connected components

0, 0, .69, .69, 1.5, 1.5, 1.8, 1.8

φ3(G ) = max

0,

2

6,

2

4

=

1

2

0, 0, .69, .69, 1.5, 1.5, 1.8, 1.8

φ3(G ) = max

0,

2

6,

2

4

=

1

2

0, .055, .055, .211, . . .

φ3(G ) = max

2

12,

2

12,

2

14

=

1

6

λk

λk = 0⇔ φk = 0

λk2≤ φk(G ) ≤ O(k2) ·

√λk

(Lee,Oveis Gharan, T, 2012)

Upper bound is constructive: in nearly-linear time we can find kdisjoint sets each of expansion at most O(k2)

√λk .

λk

λk = 0⇔ φk = 0

λk2≤ φk(G ) ≤ O(k2) ·

√λk

(Lee,Oveis Gharan, T, 2012)

Upper bound is constructive: in nearly-linear time we can find kdisjoint sets each of expansion at most O(k2)

√λk .

λk

λk2≤ φk(G ) ≤ O(k2) ·

√λk

φk(G ) ≤ O(√

(log k) · λ1.1k)

(Lee,Oveis Gharan, T, 2012)

φk(G ) ≤ O(√

(log k) · λO(k))

(Louis, Raghavendra, Tetali, Vempala, 2012)

Factor O(√

log k) is tight

λk

λk2≤ φk(G ) ≤ O(k2) ·

√λk

φk(G ) ≤ O(√

(log k) · λ1.1k)

(Lee,Oveis Gharan, T, 2012)

φk(G ) ≤ O(√

(log k) · λO(k))

(Louis, Raghavendra, Tetali, Vempala, 2012)

Factor O(√

log k) is tight

spectral clustering

The following algorithm works well in practice. (But to solve whatproblem?)

Compute k eigenvectors x(1), . . . , x(k) ofk smallest eigenvalues of L

Define mapping F : V → Rk

F (v) := (x(1)v , x

(2)v , . . . , x

(k)v )

Apply k-means to the points that vertices are mapped to

spectral embedding of a random graphk = 3

spectral clustering

LOT algorithms and LRTV algorithm find k low-conductance setsif they exist

First step is spectral embedding

Then geometric partitioning but not k-means

Cheeger inequality

λ2

2≤ φ(G ) ≤

√2λ2

φ(G ) ≤ O(k) · λ2√λk

(Kwok, Lau, Lee, Oveis Gharan, T, 2013)

and the upper bound holds for the sweep algorithm

Nearly-linear time sweep algorithm finds S such that, for every k ,

φ(S) ≤ φ(G ) · O(

k√λk

)

Cheeger inequality

λ2

2≤ φ(G ) ≤

√2λ2

φ(G ) ≤ O(k) · λ2√λk

(Kwok, Lau, Lee, Oveis Gharan, T, 2013)

and the upper bound holds for the sweep algorithm

Nearly-linear time sweep algorithm finds S such that, for every k ,

φ(S) ≤ φ(G ) · O(

k√λk

)

Cheeger inequality

λ2

2≤ φ(G ) ≤

√2λ2

φ(G ) ≤ O(k) · λ2√λk

(Kwok, Lau, Lee, Oveis Gharan, T, 2013)

and the upper bound holds for the sweep algorithm

Nearly-linear time sweep algorithm finds S such that, for every k ,

φ(S) ≤ φ(G ) · O(

k√λk

)

proof structure

1 We find a 2k-valued vector y such that

||x − y ||2 ≤ O

(λ2

λk

)where x is eigenvector of λ2

2 For every k-valued vector y , applying the sweep algorithm tox finds a set S such that

φ(S) ≤ O(k) ·(λ2 +

√λ2 · ||x − y ||

)

So the sweep algorithm finds S

φ(S) ≤ O(k) · λ2 ·1√λk

proof structure

1 We find a 2k-valued vector y such that

||x − y ||2 ≤ O

(λ2

λk

)where x is eigenvector of λ2

2 For every k-valued vector y , applying the sweep algorithm tox finds a set S such that

φ(S) ≤ O(k) ·(λ2 +

√λ2 · ||x − y ||

)

So the sweep algorithm finds S

φ(S) ≤ O(k) · λ2 ·1√λk

proof structure

1 We find a 2k-valued vector y such that

||x − y ||2 ≤ O

(λ2

λk

)where x is eigenvector of λ2

2 For every k-valued vector y , applying the sweep algorithm tox finds a set S such that

φ(S) ≤ O(k) ·(λ2 +

√λ2 · ||x − y ||

)

So the sweep algorithm finds S

φ(S) ≤ O(k) · λ2 ·1√λk

intuition for first result

• Given x , we find a 2k-valued vector y such that

||x − y ||2 ≤ O

(R(x)

λk

)

Suppose R(x) = 0 and λk > 0.

Then

• x is constant on connected components (because R(x) = 0)

• G has < k connected components (because λk > 0)

So x is at distance zero from a vector (itself) that is (k − 1)-valued

intuition for first result

• Given x , we find a 2k-valued vector y such that

||x − y ||2 ≤ O

(R(x)

λk

)

Suppose R(x) = 0 and λk > 0.

Then

• x is constant on connected components (because R(x) = 0)

• G has < k connected components (because λk > 0)

So x is at distance zero from a vector (itself) that is (k − 1)-valued

intuition for first result

• Given x , we find a 2k-valued vector y such that

||x − y ||2 ≤ O

(R(x)

λk

)

Suppose R(x) = 0 and λk > 0.

Then

• x is constant on connected components (because R(x) = 0)

• G has < k connected components (because λk > 0)

So x is at distance zero from a vector (itself) that is (k − 1)-valued

idea of proof of first result

Consider x as an embedding of vertices to the line

Partition the points vertices are mapped to into 2k intervals,minimizing 2k-means with squared-distance

“Round” to closest interval boundary

If rounded vector is far from original vector, a lot of points are farfrom boundary; then we construct witness that λk is small byfinding k disjointly supported vectors of small Rayleigh quotient

idea of proof of first result

Consider x as an embedding of vertices to the line

Partition the points vertices are mapped to into 2k intervals,minimizing 2k-means with squared-distance

“Round” to closest interval boundary

If rounded vector is far from original vector, a lot of points are farfrom boundary; then we construct witness that λk is small byfinding k disjointly supported vectors of small Rayleigh quotient

idea of proof of first result

Consider x as an embedding of vertices to the line

Partition the points vertices are mapped to into 2k intervals,minimizing 2k-means with squared-distance

“Round” to closest interval boundary

If rounded vector is far from original vector, a lot of points are farfrom boundary; then we construct witness that λk is small byfinding k disjointly supported vectors of small Rayleigh quotient

idea of proof of first result

Consider x as an embedding of vertices to the line

Partition the points vertices are mapped to into 2k intervals,minimizing 2k-means with squared-distance

“Round” to closest interval boundary

If rounded vector is far from original vector, a lot of points are farfrom boundary; then we construct witness that λk is small byfinding k disjointly supported vectors of small Rayleigh quotient

intuition for second result

• For every k-valued vector y , applying the sweep algorithm tox finds a set S such that

φ(S) ≤ O(k) ·(R(x) +

√R(x) · ||x − y ||

)

It was known:

φ(S) ≤√

2R(x)

||x − y || = 0⇒ φ(S) ≤ kR(x)

Our proof blends the two arguments

eigenvalue gap

• if λk = 0 and λk+1 > 0 then G has exactly k connectedcomponents;

• [Tanaka 2012] If λk+1 > 5k√λk , then vertices of G can be

partitioned into k sets of small conductance, each inducing asubgraph of large conductance. [Non-algorithmic proof]

• [Oveis-Gharan, T 2013] Sufficient φk+1(G ) > (1 + ε)φk(G ); ifλk+1 > O(k2)

√λk , then partition can be found

algorithmically

eigenvalue gap

• if λk = 0 and λk+1 > 0 then G has exactly k connectedcomponents;

• [Tanaka 2012] If λk+1 > 5k√λk , then vertices of G can be

partitioned into k sets of small conductance, each inducing asubgraph of large conductance. [Non-algorithmic proof]

• [Oveis-Gharan, T 2013] Sufficient φk+1(G ) > (1 + ε)φk(G ); ifλk+1 > O(k2)

√λk , then partition can be found

algorithmically

eigenvalue gap

• if λk = 0 and λk+1 > 0 then G has exactly k connectedcomponents;

• [Tanaka 2012] If λk+1 > 5k√λk , then vertices of G can be

partitioned into k sets of small conductance, each inducing asubgraph of large conductance. [Non-algorithmic proof]

• [Oveis-Gharan, T 2013] Sufficient φk+1(G ) > (1 + ε)φk(G ); ifλk+1 > O(k2)

√λk , then partition can be found

algorithmically

bounded threshold rank

[Arora, Barak, Steurer], [Barak, Raghavendra, Steurer],[Guruswami, Sinop], ...

If λk is large,then a cut of approximately minimal conductance

can be found in time exponential in kusing spectral methods [ABS]or convex relaxations [BRS, GS]

[Kwok, Lau, Lee, Oveis-Gharan, T]: the nearly-linear time sweepalgorithm finds good approximation if λk is large

[Oveis-Gharan, T]: the Linial-Goemans relaxation (solvable inpolynomial time independent of k) can be rounded better than inthe Arora-Rao-Vazirani analysis if λk is large

bounded threshold rank

[Arora, Barak, Steurer], [Barak, Raghavendra, Steurer],[Guruswami, Sinop], ...

If λk is large,then a cut of approximately minimal conductance

can be found in time exponential in kusing spectral methods [ABS]or convex relaxations [BRS, GS]

[Kwok, Lau, Lee, Oveis-Gharan, T]: the nearly-linear time sweepalgorithm finds good approximation if λk is large

[Oveis-Gharan, T]: the Linial-Goemans relaxation (solvable inpolynomial time independent of k) can be rounded better than inthe Arora-Rao-Vazirani analysis if λk is large

challenge: explain practical performanceof sweep algorithm

λ2

2≤ φ(G ) ≤ φ( output of algorithm ) ≤

√2λ2

(Kwok, Lau, Lee, Oveis Gharan, T):

φ( output of algorithm ) ≤ O(k) · λ2√λk

challenge: explain practical performanceof sweep algorithm

λ2

2≤ φ(G ) ≤ φ( output of algorithm ) ≤

√2λ2

(Kwok, Lau, Lee, Oveis Gharan, T):

φ( output of algorithm ) ≤ O(k) · λ2√λk

challenge: explain practical performanceof sweep algorithm

λ2

2≤ φ(G ) ≤ φ( output of algorithm ) ≤

√2λ2

(Spielman, Teng) in bounded-degree planar graphs:

λ2 ≤ O

(1

n

)so

φ( output of algorithm ) ≤ O

(1√n

)matching the planar separator theorem.

challenge: explain practical performanceof sweep algorithm

λ2

2≤ φ(G ) ≤ φ( output of algorithm ) ≤

√2λ2

(Spielman, Teng) in bounded-degree planar graphs:

λ2 ≤ O

(1

n

)so

φ( output of algorithm ) ≤ O

(1√n

)matching the planar separator theorem.

explain practical performance of sweepalgorithm

λ2

2≤ φ(G ) ≤ φ( output of algorithm ) ≤

√2λ2

Can we find interesting classes of graphs for which λ2 << φ(G )?

explain practical performance of sweepalgorithm

λ2

2≤ φ(G ) ≤ φ( output of algorithm ) ≤

√2λ2

Can we find interesting classes of graphs for which λ2 << φ(G )?

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